Managing storage protection faults

ABSTRACT

Management of storage used by pageable guests of a computing environment is facilitated. A query instruction is provided that details information regarding the storage location indicated in the query. It specifies whether the storage location, if protected, is protected by host-level protection or guest-level protection.

CROSS-REFERENCE TO RELATED APPLICATIONS

This application is a continuation of U.S. Ser. No. 14/570,724, entitled“USE OF TEST PROTECTION INSTRUCTION IN COMPUTING ENVIRONMENTS THATSUPPORT PAGEABLE GUESTS,” filed Dec. 15, 2014, which is a continuationof U.S. Pat. No. 8,972,670, entitled “USE OF TEST PROTECTION INSTRUCTIONIN COMPUTING ENVIRONMENTS THAT SUPPORT PAGEABLE GUESTS,” issued Mar. 3,2015, which is a continuation of U.S. Pat. No. 8,677,077, entitled “USEOF TEST PROTECTION INSTRUCTION IN COMPUTING ENVIRONMENTS THAT SUPPORTPAGEABLE GUESTS,” issued Mar. 18, 2014, which is a continuation of U.S.Pat. No. 8,364,912, entitled “USE OF TEST PROTECTION INSTRUCTION INCOMPUTING ENVIRONMENTS THAT SUPPORT PAGEABLE GUESTS,” issued Jan. 29,2013, which is a continuation of U.S. Pat. No. 8,176,280, entitled “USEOF TEST PROTECTION INSTRUCTION IN COMPUTING ENVIRONMENTS THAT SUPPORTPAGEABLE GUESTS,” issued May 8, 2012, which claims priority to U.S.Provisional Application No. 61/031,160, entitled “GUEST TEST PROTECTIONINSTRUCTION WITH EXTENDED STORAGE PROTECTION IN A VIRTUAL MACHINEENVIRONMENT,” filed Feb. 25, 2008, each of which is hereby incorporatedherein by reference in its entirety.

BACKGROUND

This invention relates, in general, to computing environments thatsupport pageable guests, and in particular, to managing use of storageby multiple pageable guests of a computing environment.

At least a portion of storage of a computing environment (also referredto herein as memory) is typically shared among a plurality of processesexecuting within the environment. This sharing of storage increases therisk of compromising the integrity of the data stored in the storage.Thus, in order to provide data integrity, certain storage protectionsare applied. These protections are used to manage the storage, andthereby, protect the data.

The protections may be applied in different types of computingenvironments, including those that support interpretative execution andpageable guests. In a pageable guest environment, for instance, multipleguests share the same storage, and again, protections are used toprovide data integrity and protect the storage. In such environments,however, information relating to the protections becomes convoluted forthe guests and/or the hosts executing the guests.

BRIEF SUMMARY

Based on the foregoing, a need exists for a capability that facilitatesmanagement of storage used by multiple pageable guests. In one example,a need exists for a capability that indicates whether a storageprotection fault was due to host level protection or guest levelprotection.

The shortcomings of the prior art are overcome and additional advantagesare provided through the provision of a computer program product forfacilitating management of storage of a computing environment. Thecomputer program product includes a computer readable storage mediumreadable by a processing circuit and storing instructions for executionby the processing circuit for performing a method. The method includesdetecting a storage protection fault in an attempt to access an area ofstorage that is protected by at least one of a host level of protectionand a guest level of protection; determining, based on one or morechecks, whether the detected storage protection fault is due to the hostlevel of protection or the guest level of protection; and performing anaction based on the determining, wherein the action performed is oneaction based on the detected storage protection fault being due to theguest level of protection and is another action based on the detectedstorage protection fault being due to the host level of protection.

Computer-implemented methods and computer systems relating to one ormore aspects of the present invention are also described and may beclaimed herein.

Additional features and advantages are realized through the techniquesof the present invention. Other embodiments and aspects of the inventionare described in detail herein and are considered a part of the claimedinvention.

BRIEF DESCRIPTION OF THE SEVERAL VIEWS OF THE DRAWINGS

One or more aspects of the present invention are particularly pointedout and distinctly claimed as examples in the claims at the conclusionof the specification. The foregoing and other objects, features, andadvantages of the invention are apparent from the following detaileddescription taken in conjunction with the accompanying drawings inwhich:

FIG. 1 depicts one embodiment of a computing environment to incorporateand use one or more aspects of the present invention;

FIG. 2 depicts one embodiment of an emulated computing environment toincorporate and use one or more aspects of the present invention;

FIG. 3 depicts one example of a frame descriptor used in accordance withan aspect of the present invention;

FIGS. 4A-4C depict examples of region table entries used in accordancewith an aspect of the present invention;

FIGS. 5A-5B depict examples of segment table entries used in accordancewith an aspect of the present invention;

FIG. 6 depicts one example of a page table entry used in accordance withan aspect of the present invention;

FIG. 7 depicts one example of a page status table entry used inaccordance with an aspect of the present invention;

FIGS. 8A-8B depict one embodiment of guest processing logic in anenvironment that supports suppression on protection and enhancedsuppression on protection facilities, in accordance with an aspect ofthe present invention;

FIG. 9 depicts one example of a format of a Test Protection instruction,in accordance with an aspect of the present invention;

FIGS. 10A-10B depict one embodiment of the logic of the Test Protectioninstruction, in accordance with an aspect of the present invention; and

FIG. 11 depicts one embodiment of a computer program productincorporating one or more aspects of the present invention.

DETAILED DESCRIPTION

In accordance with an aspect of the present invention, a capability isprovided that facilitates management of storage used by multiplepageable guests of a computing environment. As one example, an enhancedsuppression on protection facility is provided that enables thedetermination of which level of protection (host or guest) caused afault condition. As another example, a query instruction is provided(e.g., Test Protection (TPROT)) that details information regarding thearea of storage indicated in the query.

One embodiment of a computing environment to incorporate and use one ormore aspects of the present invention is described with reference toFIG. 1. Computing environment 100 is based, for instance, on thez/Architecture® offered by International Business Machines Corporation,Armonk, N.Y. The z/Architecture® is described in an IBM® publicationentitled, “z/Architecture Principles of Operation,” IBM® Publication No.SA22-7832-05, April, 2007, which is hereby incorporated herein byreference in its entirety. In one example, a computing environment basedon the z/Architecture® includes an eServer zSeries®, offered byInternational Business Machines Corporation, Armonk, N.Y. IBM®,z/Architecture® and zSeries® are registered trademarks of InternationalBusiness Machines Corporation, Armonk, N.Y., USA. Other names usedherein may be registered trademarks, trademarks, or product names ofInternational Business Machines Corporation or other companies.

As one example, computing environment 100 includes a central processorcomplex (CPC) 102 providing virtual machine support. CPC 102 includes,for instance, one or more virtual machines 104, one or more centralprocessors 106, at least one host 108 (e.g., a control program, such asa hypervisor), and an input/output subsystem 110, each of which isdescribed below. In this example, the virtual machines and host areincluded in memory.

The virtual machine support of the CPC provides the ability to operatelarge numbers of virtual machines, each capable of hosting a guestoperating system 112, such as Linux. Each virtual machine 104 is capableof functioning as a separate system. That is, each virtual machine canbe independently reset, host a guest operating system, and operate withdifferent programs. An operating system or application program runningin a virtual machine appears to have access to a full and completesystem, but in reality, only a portion of it is available.

In this particular example, the model of virtual machines is a V=Vmodel, in which the absolute or real memory of a virtual machine isbacked by host virtual memory, instead of real or absolute memory. Eachvirtual machine has a virtual linear memory space. The physicalresources are owned by host 108, and the shared physical resources aredispatched by the host to the guest operating systems, as needed, tomeet their processing demands. This V=V virtual machine (i.e., pageableguest) model assumes that the interactions between the guest operatingsystems and the physical shared machine resources are controlled by thehost, since the large number of guests typically precludes the host fromsimply partitioning and assigning the hardware resources to theconfigured guests. One or more aspects of a V=V model are furtherdescribed in an IBM® publication entitled “z/VM: Running Guest OperatingSystems,” IBM® Publication No. SC24-5997-02, October 2001, which ishereby incorporated herein by reference in its entirety.

Central processors 106 are physical processor resources that areassignable to a virtual machine. For instance, virtual machine 104includes one or more logical processors, each of which represents all ora share of a physical processor resource 106 that may be dynamicallyallocated to the virtual machine. Virtual machines 104 are managed byhost 108. As examples, the host may be implemented in microcode runningon processors 106 or be part of a host operating system executing on themachine. In one example, host 108 is a VM hypervisor, such as z/VM®,offered by International Business Machines Corporation, Armonk, N.Y. Oneembodiment of z/VM® is described in an IBM® publication entitled “z/VM:General Information Manual,” IBM Publication No. GC24-5991-04, October2001, which is hereby incorporated herein by reference in its entirety.

Input/output subsystem 110 directs the flow of information betweendevices and main storage. It is coupled to the central processingcomplex, in that it can be part of the central processing complex orseparate therefrom. The I/O subsystem relieves the central processors ofthe task of communicating directly with the I/O devices coupled to theCPC and permits data processing to proceed concurrently with I/Oprocessing.

In one embodiment, the host (e.g., z/VM®) and processor (e.g., System z)hardware/firmware interact with each other in a controlled cooperativemanner in order to process V=V guest operating system operations withoutrequiring transfer of control from/to the guest operating system and thehost. Guest operations can be executed directly without hostintervention via a facility that allows instructions to beinterpretively executed for a pageable storage mode guest. This facilityprovides an instruction, Start Interpretive Execution (SIE), which thehost can issue, designating a control block called a state descriptionwhich holds guest (virtual machine) state and controls. The instructionplaces the machine into an interpretive-execution mode in which guestinstructions and interruptions are processed directly, until a conditionrequiring host attention arises. When such a condition occurs,interpretive execution is ended, and either a host interruption ispresented, or the SIE instruction completes storing details of thecondition encountered; this latter action is called interception. Oneexample of interpretive execution is described in System/370 ExtendedArchitecture/Interpretive Execution, IBM Publication No. SA22-7095-01,September 1985, which is hereby incorporated herein by reference in itsentirety.

In one example, the interpretative execution facility is an element ofthe Processor Resource/Systems Manager® (PR/SM) offered by IBM®. Itpermits a virtual server instruction stream to be run on the processorusing a single instruction-Start Interpretive Execution (SIE). The SIEinstruction is used by the server's logical partitioning (LPAR) supportto divide, for instance, a zSeries® or S/390 processor complex intoproven secure logical partitions. (Note: The IBM® S/390 CMOS G6 familyof processors PR/SM facility received certification at the ITSEC E4level of security.)

The SIE instruction runs a virtual server dispatched by the controlprogram until the server's time slice has been consumed or until theserver wants to perform an operation that the hardware cannot virtualizeor for which the control program is to regain control. At that point,the SIE instruction ends and control is returned to the control program,which either simulates the instruction or places the virtual server inan involuntary wait state. When complete, the control program againschedules the virtual server to run, and the cycle starts again. In thisway, the full capabilities and speed of the CPU are available to thevirtual server. Only those privileged instructions that requireassistance from or validation by the control program are intercepted.These SIE intercepts, as they are known as, are also used by the controlprogram to impose limits on the operations a virtual server may performon a real device.

Moreover, this mechanism enables the control program to limit the scopeof many kinds of hardware or software failures. If the error can beisolated to a particular virtual server, only that virtual server failsand the operation can be retried or the virtual server can bereinitialized (rebooted) without affecting any testing or productionwork running in other virtual servers. The control program is designedso that failures occurring in virtual servers do not affect the controlprogram or other virtual servers.

Another example of a computing environment to incorporate one or moreaspects of the present invention is depicted in FIG. 2. In this example,an emulated host computer system 200 is provided that emulates a hostcomputer 202 of a host architecture. In emulated host computer system200, a host processor (CPU) 204 is an emulated host processor (orvirtual host processor) and is realized through an emulation processor206 having a different native instruction set architecture than used bythe processors of host computer 202. Emulated host computer system 200has memory 208 accessible to emulation processor 206. In the exampleembodiment, memory 208 is partitioned into a host computer memory 210portion and an emulation routines 212 portion. Host computer memory 210is available to programs of emulated host computer 202 according to hostcomputer architecture, and may include both a host or hypervisor 214 andone or more virtual machines 216 running guest operating systems 218,analogous to the like-named elements in FIG. 1.

Emulation processor 206 executes native instructions of an architectedinstruction set of an architecture other than that of the emulatedprocessor 204. The native instructions are obtained, for example, fromemulation routines memory 212. Emulation processor 206 may access a hostinstruction for execution from a program in host computer memory 210 byemploying one or more instruction(s) obtained in a sequence &access/decode routine which may decode the host instruction(s) accessedto determine a native instruction execution routine for emulating thefunction of the host instruction accessed. One such host instruction maybe, for example, a Start Interpretive Execution (SIE) instruction, bywhich the host seeks to execute a guest program in a virtual machine.The emulation routines 212 may include support for this instruction, andfor executing a sequence of guest instructions in a virtual machine 216in accordance with the definition of this SIE instruction.

Other facilities that are defined for host computer system 202architecture may be emulated by architected facilities routines,including such facilities as general purpose registers, controlregisters, dynamic address translation, and I/O subsystem support andprocessor cache for example. The emulation routines may also takeadvantage of functions available in emulation processor 206 (such asgeneral registers and dynamic translation of virtual addresses) toimprove performance of the emulation routines. Special hardware andoffload engines may also be provided to assist processor 206 inemulating the function of host computer 202.

Many types of computing environments, including those described above,general purpose computers, data processing systems, and others, employstorage organized using a virtual memory scheme. A virtual memory systemorganizes storage in units called blocks (or pages). These blocks aremoved between a fast, primary memory and one or more larger and usuallyslower, secondary, tertiary, etc. storage units. The movement of blocks(often called swapping) is transparent to the applications or processesthat are executed in the computing environment, enabling theapplications or processes to behave as if they each have an unlimitedamount of storage.

Although a general virtual memory allows applications and/or processesthat are executing in a computing environment to behave as if they havean unlimited amount of memory at their disposal, in actuality, theamount of storage available to a particular application or process islimited by the amount of storage in the computing environment andfurther limited by the number of concurrently executing programs sharingthat storage. A virtual memory scheme hides the actual physical addressof memory from the application programs. Application programs accesstheir memory space using a logical address (e.g., virtual address),which is then converted to a physical address by the computingenvironment.

The process of translating a virtual address during a storage referenceinto the corresponding real address or absolute address is referred toas dynamic address translation (DAT). The virtual address may be aprimary virtual address, a secondary virtual address, an Access Registerspecified virtual address, or a home virtual address. These addressesare translated by means of a primary, a secondary, an AR specified, or ahome address space control element, respectively. After selection of theappropriate address space control element, the translation process isthe same for all of the four types of virtual address. DAT may use fromfive to two levels of tables (region first table, region second table,region third table, segment table, and page table) as transformationparameters. An enhanced dynamic address translation (EDAT) process mayuse from five to one levels of table, by omitting the page table forsome or all translations. The designation (origin and length) of thehighest level table for a specific address space is called an addressspace control element, and it is found for use by DAT in a controlregister or as specified by an access register. Alternatively, theaddress space control element for an address space may be a real spacedesignation, which indicates that DAT is to translate the virtualaddress simply by treating it as a real address and without using anytables.

DAT uses, at different times, the address space control elements indifferent control registers or specified by the access registers. Thechoice is determined by the program-specified translation mode specifiedin the current PSW (Program Status Word). Four translation modes areavailable: primary space mode, secondary space mode, access registermode, and home space mode. Different address spaces are addressabledepending on the translation mode.

The result of enhanced DAT upon a virtual address may be either a realor an absolute address. If it is a real address, a prefixing operationis then performed to obtain the corresponding absolute address, whichcan be used to reference memory. Prefixing provides the ability toassign the range of real addresses 0-8191 (as an example) to a differentarea in absolute storage for each CPU, thus permitting more than one CPUsharing main storage to operate concurrently with a minimum ofinterference, especially in the processing of interruptions. Prefixingcauses real addresses in the range 0-8191 to correspond one-for-one tothe area of 8K byte absolute addresses (the prefix area) identified bythe value in bit positions 0-50 of the prefix register for the CPU, andthe area of real addresses identified by that value in the prefixregister to correspond one-for-one to absolute addresses 0-8191. Theremaining real addresses are the same as the corresponding absoluteaddresses. This transformation allows each CPU to access all of mainstorage, including the first 8K bytes and the locations designated bythe prefix registers of other CPUs.

Dynamic address translation, prefixing, and enhanced DAT are describedin more detail in U.S. Ser. No. 11/972,725, entitled, “Enhanced DynamicAddress Translation with Frame Management Function,” Gainey et al.,filed Jan. 11, 2008, which is hereby incorporated herein by reference inits entirety.

To facilitate an understanding of one or more aspects of the presentinvention, reference is made to various terms and data structures (e.g.,tables, lists), which are described below.

Frame Descriptor

A frame descriptor describes a host page frame; that is, an area of realmemory (frame) capable of holding a portion of virtual memory (page). Itis allocated, deallocated, and initialized by the host and may beupdated by Host Page Management Assist functions (as described, forinstance in U.S. Ser. No. 10/854,990, entitled “Facilitating Managementof Storage of a Pageable Mode Virtual Environment Absent Intervention ofa Host of the Environment,” Blandy et al., filed May 27, 2004, which ishereby incorporated herein by reference in its entirety).

In one example, a frame descriptor 300 (FIG. 3) is, for instance, a32-byte block residing in host home space virtual storage on a 32 byteboundary, and includes the following fields, as examples:

-   -   (a) Next Frame Descriptor Address 302: In one example, the        contents of this field, with five zeros appended on the right,        specify the host home space virtual address of the next frame        descriptor on the list. A value of zero indicates that the frame        descriptor is the last on the list.    -   This field is initialized by the host and may be changed by the        host or by Host Page Management Assist functions.    -   (b) Page Frame Real Address or PTE Copy 304: When the frame        descriptor is in the available frame descriptor list (AFDL), the        contents of this field, with twelve zeros appended on the right,        specify the host real address of the first byte (byte 0) of a        host frame that is available for allocation to provide host        storage.    -   When the frame descriptor is in a processed frame descriptor        list (PFDL), this field includes a copy of the page table entry        (PTE) designated by the page table entry address field, as it        appeared before the host page was resolved.    -   This field is initialized by the host and may be changed by the        host or by Host Page Management Assist functions.    -   (c) Page Table Entry Address 306: When the frame descriptor is        on the processed frame descriptor list, the contents of this        field, with three zeros appended on the right, specify the host        real or host absolute address of the page table entry for the        host virtual page.    -   This field is initialized by the host and may be changed by the        host or by Host Page Management Assist functions.

Multiple frame descriptors may be linked to one another to form a list,such as an available frame descriptor list (AFDL) or a processed framedescriptor list (PFDL). A frame descriptor exists in one of the twolists. A separate pair of these lists is provided for each CPU. Theorigins of the AFDL and PFDL for a CPU are designated by means of fieldsin the prefix area of the CPU.

The available frame descriptor list (AFDL) is a list of framedescriptors that describes host frames the host has cleared and has madeavailable for allocation to host pages. The AFDL is designated by anAFDL origin (AFDLO) at a specified host real address.

The contents of the AFDLO, with five zeros appended on the right,specify the host home space virtual address of the first framedescriptor on the AFDL. A value of zero indicates that the list isempty.

The AFDLO is initialized by the host and may be changed by the host orHost Page Management Assist functions. The AFDLO is changed, in oneembodiment, by means of a non-interlocked update operation.

The processed frame descriptor list (PFDL) is a list of framedescriptors that describes host frames that have been used to resolvehost page invalid conditions during guest interpretation. The hostframes that are described by the PFDL have been assigned to host pagesthat provide storage for a guest. The PFDL is designated by a PFDLorigin (PFDLO) at a specified host real address. The contents of thePFDLO, with five zeros appended on the right, specify the host homespace virtual address of the first frame descriptor on the PFDL. A valueof zero indicates that the list is empty.

The PFDLO is initialized by the host and may be changed by the host or aHost Page Management Assist function. The PFDLO is changed, in oneembodiment, by means of a doubleword concurrent interlocked updateoperation that maintains the integrity of the list.

Region Table Entries

The term “region table entry” indicates a region first table entry, aregion second table entry, or a region third table entry. The level(first, second, or third) of the table containing an entry is identifiedby the table type (TT) bits in the entry. Examples of the formats ofentries fetched from the region first table, region second table, andregion third table are depicted in FIGS. 4A-4C. In particular, FIG. 4Adepicts one embodiment of the format of a Region First Table entry 400;FIG. 4B depicts one embodiment of the format of a Region Second Tableentry 430; and FIG. 4C depicts one embodiment of the format of a RegionThird Table entry 460.

As examples, the fields in the three levels of region table entries areallocated as follows:

Region Second Table Origin 402, Region Third Table Origin 432, andSegment Table Origin 462: A region first table entry includes a regionsecond table origin. A region second table entry includes a region thirdtable origin. A region third table entry includes a segment tableorigin. The following description applies to each of the three origins.In one example, bits 0-51 of the entry, with 12 zeros appended on theright, form a 64-bit address that designates the beginning of the nextlower level table.

DAT Protection Bit (P) 406, 436, 466: When enhanced DAT applies, bit 54is treated as being OR'ed with the DAT protection bit in each subsequentregion table entry, segment table entry, and, when applicable, pagetable entry used in the translation. Thus, when the bit is one, DATprotection applies to the entire region or regions specified by theregion table entry. When the enhanced DAT facility is not installed, orwhen the facility is installed but the enhanced DAT enablement controlis zero, bit 54 of the region table entry is ignored.

Region Second Table Offset 408, Region Third Table Offset 438, andSegment Table Offset (TF) 468: A region first table entry includes aregion second table offset. A region second table entry includes aregion third table offset. A region third table entry includes a segmenttable offset. The following description applies to each of the threeoffsets. Bits 56 and 57 of the entry specify the length of a portion ofthe next lower level table that is missing at the beginning of thetable; that is, the bits specify the location of the first entryactually existing in the next lower level table. The bits specify thelength of the missing portion in units of 4,096 bytes, thus making thelength of the missing portion variable in multiples of 512 entries. Thelength of the missing portion, in units of 4,096 bytes, is equal to theTF value. The contents of the offset field, in conjunction with thelength field, bits 62 and 63, are used to establish whether the portionof the virtual address (RSX, RTX, or SX) to be translated by means ofthe next lower level table designates an entry that actually exists inthe table.

Region Invalid Bit (I) 410, 440, 470: A region is a contiguous range of,for example, 2 gigabytes of virtual addresses. Bit 58 in a region firsttable entry or region second table entry controls whether the set ofregions associated with the entry is available. Bit 58 in a region thirdtable entry controls whether the single region associated with the entryis available. When bit 58 is zero, address translation proceeds by usingthe region table entry. When the bit is one, the entry cannot be usedfor translation.

Table Type Bits (TT) 412, 442, 472: Bits 60 and 61 of the region firsttable entry, region second table entry, and region third table entryidentify the level of the table containing the entry, as follows: Bits60 and 61 identify the correct table level, considering the type oftable designation that is the address space control element being usedin the translation and the number of table levels that have so far beenused; otherwise, a translation specification exception is recognized. Asan example, the following table shows the table type bits:

Table Type bits for region table Entries Bits 60 and 61 Region-TableLevel 11 First 10 Second 01 Third

Region Second Table Length 414, Region Third Table Length 444, andSegment Table Length 474 (TL): A region first table entry includes aregion second table length. A region second table entry includes aregion third table length. A region third table entry includes a segmenttable length. The following description applies to each of the threelengths. Bits 62 and 63 of the entry specify the length of the nextlower level table in units of 4,096 bytes, thus making the length of thetable variable in multiples of 512 entries. The length of the next lowerlevel table, in units of 4,096 bytes, is one more than the TL value. Thecontents of the length field, in conjunction with the offset field, bits56 and 57, are used to establish whether the portion of the virtualaddress (RSX, RTX, or SX) to be translated by means of the next lowerlevel table designates an entry that actually exists in the table.

All other bit positions of the region table entry are reserved forpossible future extensions and should contain zeros; otherwise, theprogram may not operate compatibly in the future. When enhanced DATapplies, the reserved bit positions of the region table entry shouldcontain zeros even if the table entry is invalid.

Segment Table Entries

When enhanced DAT does not apply, or when enhanced DAT applies and theSTE format control, bit 53 of the segment table entry, is zero, theentry fetched from the segment table has the format (e.g., Format 0) asdepicted in FIG. 5A. When enhanced DAT applies and the STE formatcontrol is one, the entry fetched from the segment table has, forexample, the format (e.g., Format 1) as depicted in FIG. 5B.

As one example, a Format 0 segment table entry 500 (FIG. 5A) includesthe following fields:

Page Table Origin 502: When enhanced DAT does not apply, or whenenhanced DAT applies but the STE format control, bit 53 of the segmenttable entry, is zero, bits 0-52, with 11 zeros appended on the right,form a 64-bit address that designates the beginning of a page table. Itis unpredictable whether the address is real or absolute.

STE Format Control (FC) 506: When enhanced DAT applies, bit 53 is theformat control for the segment table entry, as follows:

-   -   When the FC bit is zero, bits 0-52 of the entry form the page        table origin, and bit 55 is reserved.    -   When the FC bit is one, bits 0-43 of the entry form the segment        frame absolute address, bit 47 is the ACCF validity control,        bits 48-51 are the access control bits, bit 52 is the fetch        protection bit, and bit 55 is the change recording override.        When enhanced DAT does not apply, bit 53 is ignored.

DAT Protection Bit (P) 508: Bit 54, when one, indicates that DATprotection applies to the entire segment.

-   -   When enhanced DAT does not apply, bit 54 is treated as being        OR'ed with the DAT protection bit in the page table entry used        in the translation.    -   When enhanced DAT applies, the DAT protection bit in any and all        region table entries used in the translation are treated as        being OR'ed with the DAT protection bit in the segment table        entry; when the STE format control is zero, the DAT protection        bit in the STE is further treated as being OR'ed with the DAT        protection bit in the page table entry.

Segment Invalid Bit (I) 510: Bit 58 controls whether the segmentassociated with the segment table entry is available.

-   -   When the bit is zero, address translation proceeds by using the        segment table entry.    -   When the bit is one, the segment table entry cannot be used for        translation.

Common Segment Bit (C) 512: Bit 59 controls the use of the translationlookaside buffer (TLB) copies of the segment table entry. When enhancedDAT does not apply or when enhanced DAT applies but the format controlis zero, bit 59 also controls the use of the TLB copies of the pagetable designated by the segment table entry.

-   -   A zero identifies a private segment; in this case, the segment        table entry and any page table it designates may be used only in        association with the segment table origin that designates the        segment table in which the segment table entry resides.    -   A one identifies a common segment; in this case, the segment        table entry and any page table it designates may continue to be        used for translating addresses corresponding to the segment        index, even though a different segment table is specified.

However, TLB copies of the segment table entry and any page table for acommon segment are not usable if the private space control, bit 55, isone in the address space control element used in the translation or ifthat address space control element is a real space designation. Thecommon segment bit is to be zero if the segment table entry is fetchedfrom storage during a translation when the private space control is onein the address space control element being used; otherwise, atranslation specification exception is recognized.

Table Type Bits (TT) 514: Bits 60 and 61 of the segment table entry are00 binary to identify the level of the table containing the entry. Themeanings of possible values of bits 60 and 61 in a region table entry orsegment table entry are as follows:

Table Type Bits 60, 61 Bits 60 and 61 Table Level 11 Region-first 10Region-second 01 Region-third 00 Segment

Bits 60 and 61 are to identify the correct table level, considering thetype of table designation that is the address space control elementbeing used in the translation and the number of table levels that haveso far been used; otherwise, a translation specification exception isrecognized.

All other bit positions of the segment table entry are reserved forpossible future extensions and should contain zeros; otherwise, theprogram may not operate compatibly in the future. When enhanced DATapplies, the reserved bit positions of the segment table entry shouldcontain zeros even if the table entry is invalid.

As one example, a Format 1 segment table entry 550 (FIG. 5B) includesthe following fields:

Segment Frame Absolute Address (SFAA) 552: When enhanced DAT applies andthe STE format control is one, bits 0-43 of the entry, with 20 zerosappended on the right, form the 64-bit absolute address of the segment.

ACCF Validity Control (AV) 556: When enhanced DAT applies and the STEformat control is one, bit 47 is the access control bits and fetchprotection bit (ACCF) validity control. When the AV control is zero,bits 48-52 of the segment table entry are ignored. When the AV controlis one, bits 48-52 are used as described below.

Access Control Bits (ACC) 558: When enhanced DAT applies, the STE formatcontrol is one, and the AV control is one, bits 48-51 of the segmenttable entry include the access control bits that may be used for any keycontrolled access checking that applies to the address.

Fetch Protection Bit (F) 560: When enhanced DAT applies, the STE formatcontrol is one, and the AV control is one, bit 52 of the segment tableentry includes the fetch protection bit that may be used for any keycontrolled access checking that applies to the address.

STE Format Control (FC) 562: When enhanced DAT applies, bit 53 is theformat control for the segment table entry, as follows:

-   -   When the FC bit is zero, bits 0-52 of the entry form the page        table origin, and bit 55 is reserved.    -   When the FC bit is one, bits 0-43 of the entry form the segment        frame absolute address, bit 47 is the ACCF validity control,        bits 48-51 are the access control bits, bit 52 is the fetch        protection bit, and bit 55 is the change recording override.        When enhanced DAT does not apply, bit 53 is ignored.

DAT Protection Bit (P) 564: Bit 54, when one, indicates that DATprotection applies to the entire segment.

-   -   When enhanced DAT does not apply, bit 54 is treated as being        OR'ed with the DAT protection bit in the page table entry used        in the translation.    -   When enhanced DAT applies, the DAT protection bit in any and all        region table entries used in the translation are treated as        being OR'ed with the DAT protection bit in the segment table        entry; when the STE format control is zero, the DAT protection        bit in the STE is further treated as being OR'ed with the DAT        protection bit in the page table entry.

Change Recording Override (CO) 566: When enhanced DAT applies, and theSTE format control is one, bit 55 of the segment table entry is thechange recording override for the segment. When enhanced DAT does notapply, or when enhanced DAT applies but the STE format control is zero,bit 55 of the segment table entry is ignored.

Segment Invalid Bit (I) 568: Bit 58 controls whether the segmentassociated with the segment table entry is available.

-   -   When the bit is zero, address translation proceeds by using the        segment table entry.    -   When the bit is one, the segment table entry cannot be used for        translation.

Common Segment Bit (C) 570: Bit 59 controls the use of the translationlookaside buffer (TLB) copies of the segment table entry. When enhancedDAT does not apply or when enhanced DAT applies but the format controlis zero, bit 59 also controls the use of the TLB copies of the pagetable designated by the segment table entry.

-   -   A zero identifies a private segment; in this case, the segment        table entry and any page table it designates may be used only in        association with the segment table origin that designates the        segment table in which the segment table entry resides.    -   A one identifies a common segment; in this case, the segment        table entry and any page table it designates may continue to be        used for translating addresses corresponding to the segment        index, even though a different segment table is specified.

However, TLB copies of the segment table entry and any page table for acommon segment are not usable if the private space control, bit 55, isone in the address space control element used in the translation or ifthat address space control element is a real space designation. Thecommon segment bit is to be zero if the segment table entry is fetchedfrom storage during a translation when the private space control is onein the address space control element being used; otherwise, atranslation specification exception is recognized.

Table Type Bits (TT) 572: Bits 60 and 61 of the segment table entry are00 binary to identify the level of the table containing the entry. Themeanings of possible values of bits 60 and 61 in a region table entry orsegment table entry are as follows:

Table Type Bits 60, 61 Bits 60 and 61 Table Level 11 Region-first 10Region-second 01 Region-third 00 Segment

Bits 60 and 61 are to identify the correct table level, considering thetype of table designation that is the address space control elementbeing used in the translation and the number of table levels that haveso far been used; otherwise, a translation specification exception isrecognized.

All other bit positions of the segment table entry are reserved forpossible future extensions and should contain zeros; otherwise, theprogram may not operate compatibly in the future. When enhanced DATapplies, the reserved bit positions of the segment table entry shouldcontain zeros even if the table entry is invalid.

Page Table Entries

The state information for guest blocks (e.g., an area (e.g., 4 K-Bytes)in absolute memory that has associated therewith a single storage keyand CMM state) is maintained, for instance, in host page tables (PTs)and page status tables (PGSTs) that describe a guest's memory. Thesetables include, for instance, one or more page table entries (PTEs) andone or more page status table entries (PGSTEs), respectively, which aredescribed in further detail below.

One example of a page table entry 600 is described with reference toFIG. 6. In one embodiment, the fields in the page table entry areallocated as follows:

Page Frame Real Address (PFRA) 602: Bits 0-51 provide the leftmost bitsof a real (in this case host real) storage address. When these bits areconcatenated with the 12-bit byte index field of the virtual address onthe right, a 64-bit real address is obtained.

Page Invalid Bit (I) 604: Bit 53 controls whether the page associatedwith the page table entry is available. When the bit is zero, addresstranslation proceeds by using the page table entry. Further, with regardto collaborative memory management (CMM) between host and guest, thehost state is r (resident; i.e., the guest block is present in a hostframe). When the bit is one, the page table entry is not used fortranslation, and the CMM host state is p (preserved; i.e., the guestblock is not present in a host frame, but has been preserved by the hostin some auxiliary storage) or z (logically zero; i.e., the guest blockis not present in a host frame and the contents of the guest block areknown to be zeros), as determined by PGSTE.Z.

DAT Protection Bit (P) 606: Bit 54 controls whether store accesses canbe made in the page. This protection mechanism is in addition to the keycontrolled protection and low address protection mechanisms. The bit hasno effect on fetch accesses. If the bit is zero, stores are permitted tothe page, subject to the following additional constraints:

-   -   The DAT protection bit being zero in the segment table entry        used in the translation.    -   When enhanced DAT applies, the DAT protection bit being zero in        all region table entries used in the translation.

If the DAT protection bit is one, stores are disallowed. When no higherpriority exception conditions exist, an attempt to store when the DATprotection bit is one causes a protection exception to be recognized.The DAT protection bit in the segment table entry is treated as beingOR'ed with bit 54 when determining whether DAT protection applies to thepage. When enhanced DAT applies, the DAT protection bit in any regiontable entries used in translation are also treated as being OR'ed withbit 54 when determining whether DAT protection applies.

Other protection mechanisms, such as key-controlled protection,low-address protection, and access-list-controlled protection, may applyindependently of DAT protection and may also prohibit accesses.

Change Recording Override (CO) 608: When enhanced DAT does not apply,bit 55 of the page table entry is to contain zero; otherwise, atranslation specification exception is recognized as part of theexecution of an instruction using that entry for address translation.When enhanced DAT applies and the STE format control is zero, bit 55 ofthe page table entry is the change recording override for the page.

In addition to the above, in one example, bit position 52 of the entryis to contain zero; otherwise, a translation specification exception isrecognized as part of the execution of an instruction using that entryfor address translation. Bit positions 56-63 are not assigned and areignored.

One example of a page status table entry is described with reference toFIG. 7. A page status table entry 700 includes, for instance, thefollowing:

-   -   (a) Acc 702: Access control key;    -   (b) FP 704: Fetch protection indicator;    -   (c) Page Control Interlock (PCL) 706: This is the interlock        control for serializing updates to a page table entry (PTE) and        corresponding PGSTE, except for the PGSTE status area and PGSTE        bits that are marked as reserved.    -   (d) HR 708: Host reference backup indicator;    -   (e) HC 710: Host change backup indicator;    -   (f) GR 712: Guest reference backup indicator;    -   (g) GC 714: Guest change backup indicator;    -   (h) Status 716: Intended for host program use.    -   (i) Page Content Logically Zero Indicator (Z) 718: This bit is        meaningful when the corresponding PTE page invalid (PTE.I) bit        is one.    -   When Z is one, the content of the page that is described by this        PGSTE and corresponding PTE is considered to be zero. Any prior        content of the page does not have to be preserved and may be        replaced by a page of zeros.    -   When Z is zero, the content of the page described by the PGSTE        and corresponding PTE is not considered to be zero. The content        of the page is preserved by the host.    -   When the Z bit is one and the corresponding PTE.I bit is one,        the CMM host state is z (logically zero). This means that the        page content may be replaced by the host or by a function of the        Host Page Management Assist facility.    -   When the Z bit is one, the corresponding PTE.I bit is one, and        the page content is replaced, the page should be replaced by        associating it with a frame that has been set to zeros.    -   When the Z bit is zero and the PTE invalid bit is one, the CMM        host state is p (preserved).    -   (j) Page Class (PC) 720: When zero, the page described by the        PGSTE and corresponding PTE is a class 0 page and the delta        pinned page count array (DPPCA) for class 0 pages is used for        counting pinning and unpinning operations for the page. When        one, the page described by the PGSTE and corresponding PTE is a        class 1 page and the DPPCA for class 1 pages is used for        counting pinning and unpinning operations for the page.    -   (k) Pin Count Overflow (PCO) 722: When one, the pin count field        is in an overflow state. In this case, the total pin count is        kept by the host in another data structure not accessed by the        machine. When zero, the pin count field is not in an overflow        state.    -   (l) Frame Descriptor On Processed Frame Descriptor List (FPL)        724: When one, a frame descriptor for the page described by the        PGSTE and corresponding PTE is in a processed frame descriptor        list. The frame descriptor identifies the host frame that was        used by a HPMA resolve host page function for the page.    -   (m) Page Content Replacement Requested (PCR) 726: When one, page        content replacement was requested when the HPMA resolve host        page function was invoked for the page represented by the PGSTE        and corresponding PTE.    -   (n) Usage State (US) 728: For collaborative memory management        between host and guest, this field indicates whether the guest        state is S (stable; i.e., the contents of a stable block remain        equal to what was set by the guest); U (unused; i.e., the        contents of an unused block are not meaningful to the guest); V        (volatile; i.e., the contents of a volatile block are meaningful        to the guest, but the host may at any time discard the contents        of the block and reclaim the backing frame); or P (potentially        volatile; i.e., the contents of a potentially volatile block are        meaningful to the guest, but based upon guest change history,        the host either may discard or should preserve the contents of        the block).    -   (o) Status 730: Intended for host program use.    -   (p) Pin Count 732: An unsigned binary integer count used to        indicate whether the content of the host virtual page        represented by the PGSTE and corresponding PTE is pinned in the        real host frame specified by the page frame real address field        of the PTE. When the value of this field is greater than zero or        the page count overflow (PCO) bit is one, the corresponding page        is considered to be pinned. When the value of this field is zero        and the PCO bit is zero, the corresponding page is not        considered to be pinned.    -   At the time a page is pinned by either the host or the CPU, this        field should be incremented by 1. At the time a page is unpinned        by either the host or the CPU, this field should be decremented        by 1.    -   When the value of the pin count field is greater than zero or        the PCO bit is one, the corresponding PTE.I (page invalid) bit        is to be zero. Otherwise, unpredictable results may occur.    -   While a page is pinned, the host program should not change the        contents of the PTE page frame real address (PFRA) field, the        setting of the PTE page invalid (I) bit, or the setting of the        page protection (P) bit in the PTE or segment table entry (STE).        Otherwise unpredictable results may occur.

Further details regarding page table entries and page tables, as well assegment table entries, are provided in an IBM® publication entitled,“z/Architecture Principles of Operation,” IBM® Publication No.SA22-7832-05, April 2007, which is hereby incorporated herein byreference in its entirety. Moreover, further details regarding the PGSTEare described in U.S. Ser. No. 10/854,990, entitled “FacilitatingManagement of Storage of a Pageable Mode Virtual Environment AbsentIntervention of a Host of the Environment,” Blandy et al., filed May 27,2004; and in U.S. Patent Application Publication No. US 2007/0016904 A1,entitled, “Facilitating Processing Within Computing EnvironmentSupporting Pageable Guests,” Adlung et al., published Jan. 18, 2007,each of which is hereby incorporated herein by reference in itsentirety.

In one embodiment, there is one page status table per page table, thepage status table is the same size as the page table, a page statustable entry is the same size as a page table entry, and the page statustable is located at a fixed displacement (in host real memory) from thepage table. Thus, there is a one-to-one correspondence between each pagetable entry and page status table entry. Given the host's virtualaddress of a page, both the machine and the host can easily locate thepage status table entry that corresponds to a page table entry for aguest block.

Copy-on-Write

At times, portions of memory are to be copied. This copying can eitherbe user-initiated or initiated by an operating system. Conventionalsystems often use a lazy copy technique for a flash copy, in which thestorage to be copied is assigned a status of read-only, but the actualcopy is deferred until later. If an attempt is made to write into eitherthe original or the copy, then the memory is copied at that time andboth the original and the copy are given an input/output (I/O) status ofread-write. In this way, it appears that a copy was made immediately,but the actual copying is deferred until later (e.g., the latestpossible time). If no write is performed, no copying occurs. For thisreason, this method is called copy-on-write (COW) or virtual copy.

Generally, a copy-on-write operation is computationally expensivebecause a single write results in two write operations. That is, anexisting data block is copied from an old physical block to a newphysical block, and then the actual update/write operation is performedon the new physical block.

Instruction Execution

In the z/Architecture®, for example, instruction execution ends in oneof five ways: completion, nullification, suppression, termination, andpartial completion, each of which is described below.

Completion of instruction execution provides results as called for inthe definition of the instruction. When an interruption occurs after thecompletion of the execution of an instruction, the instruction addressin the old PSW designates the next sequential instruction.

Nullification of instruction execution has the same effect assuppression, described below, except that when an interruption occursafter the execution of an instruction has been nullified, theinstruction address in the old PSW designates the instruction whoseexecution was nullified (or an Execute instruction, as appropriate),instead of the next sequential instruction.

Suppression of instruction execution causes the instruction to beexecuted as if it specified “no operation.” The contents of any resultfields, including the condition code, are not changed. The instructionaddress in the old PSW on an interruption after suppression designatesthe next sequential instruction.

Termination of instruction execution causes the contents of any fieldsdue to be changed by the instruction to be unpredictable. The operationmay replace all, part, or none of the contents of the designated resultfields and may change the condition code if such change is called for bythe instruction. Unless the interruption is caused by a machine-checkcondition, the validity of the instruction address in the PSW, theinterruption code, and the instruction length code are not affected, andthe state or the operation of the machine is not affected in any otherway. The instruction address in the old PSW on an interruption aftertermination designates the next sequential instruction.

Partial completion of instruction execution occurs for interruptibleinstructions.

Although the execution of an instruction is treated as a no-operationwhen suppression or nullification occurs, stores may be performed as theresult of the implicit tracing action associated with some instructions.

Protection

Four protection facilities are provided to protect the contents of mainstorage from destruction or misuse by programs that contain errors orare unauthorized: key-controlled protection, access-list-controlledprotection, page protection, and low-address protection. The protectionfacilities are applied independently; access to main storage ispermitted when none of the facilities prohibits the access. Eachprotection facility is described below.

Key-Controlled Protection: Key-controlled protection affords protectionagainst improper storing or against both improper storing and fetching,but not against improper fetching alone.

When key-controlled protection applies to a storage access, a store ispermitted when the storage key matches the access key associated withthe request for storage access; a fetch is permitted when the keys matchor when the fetch-protection bit of the storage key is zero. The keysare said to match when the four access control bits of the storage keyare equal to the access key, or when the access key is zero. Theprotection action is summarized in the following table:

Is Access to Conditions Storage Fetch-Protection Bit Permitted ofStorage Key Key Relation Fetch Store 0 Match Yes Yes 0 Mismatch Yes No 1Match Yes Yes 1 Mismatch No No Explanation: Match The fouraccess-control bits of the storage key are equal to the access key, orthe access key is zero. Yes Access is permitted. No Access is notpermitted. On fetching, the information is not made available to theprogram; on storing, the contents of the storage location are notchanged.

When the access to storage is initiated by the CPU and key-controlledprotection applies, the PSW key is the access key, except that theaccess key is specified in a general register for the first operand ofcertain instructions, such as Move To Secondary and Move WithDestination Key; for the second operand of instructions, such as Move ToPrimary, Move With Key, and Move With Source Key; and for either thefirst or the second operand of, for instance, a Move Page instruction.The PSW key occupies, for instance, bit positions 8-11 of the currentPSW. When the access to storage is for the purpose of channel-programexecution, the sub-channel key associated with that channel program isthe access key. The sub-channel key for a channel program is specifiedin the operation-request block (ORB). When, for purposes ofchannel-subsystem monitoring, an access to the measurement block ismade, the measurement-block key is the access key. The measurement-blockkey is specified by, for instance, a Set Channel Monitor instruction.

When a CPU access is prohibited because of key controlled protection,the execution of the instruction is terminated, and a programinterruption for a protection exception takes place. However, the unitof operation or the execution of the instruction may be suppressed. Whena channel program access is prohibited, the start function is ended, andthe protection-check condition is indicated in the associatedinterruption-response block (IRB). When a measurement-block access isprohibited, the I/O measurement-block protection-check condition isindicated. When a store access is prohibited because of key controlledprotection, the contents of the protected location remain unchanged.When a fetch access is prohibited, the protected information is notloaded into a register, moved to another storage location, or providedto an I/O device. For a prohibited instruction fetch, the instruction issuppressed, and an arbitrary instruction-length code is indicated.

Key-controlled protection is independent of whether the CPU is in theproblem or the supervisor state and, with some exceptions, does notdepend on the type of CPU instruction or channel-command word beingexecuted. Except where otherwise specified, accesses to storagelocations that are explicitly designated by the program and that areused by the CPU to store or fetch information are subject tokey-controlled protection. Key-controlled protection does not apply whenthe storage-protection-override control is one and the value of the fouraccess-control bits of the storage key is, for instance, 9.Key-controlled protection for fetches may or may not apply when thefetch-protection-override control is one, depending on the effectiveaddress and the private-space control. The storage-protection-overridecontrol and fetch protection-override control do not affect storagereferences made by the channel subsystem. Accesses to the second operandof, for instance, a Test Block instruction are not subject tokey-controlled protection. Storage accesses by the channel subsystem toaccess the I/O measurement block, or by a channel program to fetch aCCW, IDAW (indirect data address word), or MIDAW (modified indirect dataaddress word) or to access a data area designated during the executionof a CCW, are subject to key-controlled protection. However, if a CCW,an IDAW, a MIDAW, or output data is prefetched, a protection check isnot indicated until the CCW, IDAW, or MIDAW is due to take control oruntil the data is due to be written. Key-controlled protection is notapplied to accesses that are implicitly made for any of such sequencesas:

-   -   An interruption    -   CPU logout    -   Fetching of table entries for access-register translation,        dynamic-address translation, PCnumber translation, ASN        translation, or ASN authorization    -   Tracing    -   A store-status function    -   Storing in real locations 184-191 when a Test Pending        Interruption has an operand address of zero    -   Initial program loading.

Similarly, protection does not apply to accesses initiated via theoperator facilities for altering or displaying information. However,when the program explicitly designates these locations, they are subjectto protection.

Access which are subject to key-controlled protection may be effected bya storage-protection-override control. As an example, bit 39 of controlregister 0 is the storage-protection-override control. When this bit isone, storage-protection override is active. When this bit is zero,storage protection override is inactive. When storage-protectionoverride is active, key-controlled storage protection is ignored forstorage locations having an associated storage-key value of, forinstance, 9. When storage protection override is inactive, no specialaction is taken for a storage-key value of 9. Storage-protectionoverride applies to instruction fetch and to the fetch and storeaccesses of instructions whose operand addresses are logical, virtual,or real. It does not apply to accesses made for the purpose ofchannel-program execution or for the purpose of channel-subsystemmonitoring. Storage-protection override has no effect on accesses whichare not subject to key-controlled protection.

Storage-protection override can be used to improve reliability in thecase when a possibly erroneous application program is executed inconjunction with a reliable subsystem, provided that the applicationprogram needs to access only a portion of the storage accessed by thesubsystem. The technique for doing this is as follows. The storageaccessed by the application program is given storage key 9. The storageaccessed by only the subsystem is given some other nonzero storage key,for example, key 8. The application is executed with PSW key 9. Thesubsystem is executed with PSW key 8 (in this example). As a result, thesubsystem can access both the key-8 and the key-9 storage, while theapplication program can access only the key-9 storage.

Storage-protection override affects the accesses to storage made by theCPU and also affects the result set by Test Protection. However, thoseinstructions which, in the problem state, test the PSW-key mask todetermine if a particular key value may be used are not affected bywhether storage-protection override is active. These instructionsinclude, among others, Move With Key and Set PSW Key From Address. Topermit these instructions to use an access key of 9 in the problemstate, bit 9 of the PSW-key mask is to be one.

In addition to the storage-protection-override control, accesses subjectto key-controlled protection may be effected by afetch-protection-override control. As an example, bit 38 of controlregister 0 is the fetch-protection-override control. When the bit isone, fetch protection is ignored for locations at, for instance,effective addresses 0-2047. An effective address is the address whichexists before any transformation by dynamic address translation orprefixing. However, fetch protection is not ignored if the effectiveaddress is subject to dynamic address translation and the private-spacecontrol, e.g., bit 55, is one in the address-space-control element usedin the translation. Fetch-protection override applies to instructionfetch and to the fetch accesses of instructions whose operand addressesare logical, virtual, or real. It does not apply to fetch accesses madefor the purpose of channel-program execution or for the purpose ofchannel-subsystem monitoring. When this bit is set to zero, fetchprotection of locations at effective addresses 0-2047 is determined bythe state of the fetch-protection bit of the storage key associated withthose locations. Fetch-protection override has no effect on accesseswhich are not subject to key-controlled protection.

The fetch-protection-override control allows fetch protection oflocations at, for instance, addresses 2048-4095 along with no fetchprotection of locations at addresses 0-2047.

Access-List-Controlled Protection

In the access-register mode, bit 6 of the access-list entry, thefetch-only bit, controls which types of operand references are permittedto the address space specified by the access-list entry. When the entryis used in the access-register-translation part of a reference and bit 6is zero, both fetch-type and store-type references are permitted; whenbit 6 is one, fetch-type references are permitted, and an attempt tostore causes a protection exception to be recognized and the executionof the instruction to be suppressed. The fetch-only bit is included inthe ALB access-list entry. A change to the fetch-only bit in anaccess-list entry in main storage does not necessarily have animmediate, if any, effect on whether a protection exception isrecognized. However, this change to the bit does have an effectimmediately after, for example, Purge ALB or a Compare And Swap AndPurge instruction that purges the ALB is executed. A Test Protectioninstruction, described below, takes into considerationaccess-list-controlled protection when the CPU is in the access-registermode. A violation of access-list controlled protection causes conditioncode 1, as an example, to be set, except that it does not prevent, forinstance, condition code 2 or 3 from being set when the conditions forthose codes are satisfied.

A violation of access-list-controlled protection causes suppression. Aviolation of any of the other protection types may cause termination.

Page Protection

The page-protection facility controls access to virtual storage by usingthe page-protection bit in each page-table entry and segment-tableentry. It provides protection against improper storing. Thepage-protection bit, bit 54 of the page-table entry, controls whetherstoring is allowed into the corresponding 4K-byte page. When the bit iszero, both fetching and storing are permitted; when the bit is one, onlyfetching is permitted. When an attempt is made to store into a protectedpage, the contents of the page remain unchanged, the unit of operationor the execution of the instruction is suppressed, and a programinterruption for protection takes place. The page-protection bit, bit 54of the segment-table entry, is treated as being ORed into thepage-protection-bit position of each entry in the page table designatedby the segment-table entry. Thus, when the segment-table-entrypage-protection bit is one, the effect is as if the page-protection bitwere one in each entry in the designated page table. Page protectionapplies to store-type references that use a virtual address.

Low-Address Protection

The low-address-protection facility provides protection against thedestruction of main-storage information used by the CPU duringinterruption processing. This is accomplished by prohibitinginstructions from storing with effective addresses in, for instance, theranges 0 through 511 and 4096 through 4607 (the first 512 bytes of eachof the first and second 4K-byte effective-address blocks). The rangecriterion is applied before address transformation, if any, of theaddress by dynamic address translation or prefixing. However, the rangecriterion is not applied, with the result that low-address protectiondoes not apply, if the effective address is subject to dynamic addresstranslation and the private-space control, bit 55, is one in theaddress-space-control element used in the translation. Low-addressprotection does not apply if the address-space-control element to beused is not available due to another type of exception. Low-addressprotection is under control of bit 35 of control register 0, thelow-address-protection-control bit. When the bit is zero, low-addressprotection is off; when the bit is one, low-address protection is on.

If an access is prohibited because of low-address protection, thecontents of the protected location remain unchanged, the execution ofthe instruction is terminated, and a program interruption for aprotection exception takes place. However, the unit of operation or theexecution of the instruction may be suppressed. An attempt by theprogram to store by using effective addresses in, for instance, therange 0 through 511 or 4096 through 4607 is subject to low-addressprotection. Low-address protection is applied to the store accesses ofinstructions whose operand addresses are logical, virtual, or real.Low-address protection is also applied to the trace table. Low-addressprotection is not applied to accesses made by the CPU or the channelsubsystem for such sequences as interruptions, CPU logout, the storingof the I/O-interruption code in, for instance, real locations 184-191 byTest Pending Interruption, and the initial-program-loading andstore-status functions, nor is it applied to data stores during I/O datatransfer. However, explicit stores by a program at any of theselocations are subject to low-address protection.

Low-address protection and key-controlled protection apply to the samestore accesses, except that: low-address protection does not apply tostoring performed by the channel subsystem, whereas key-controlledprotection does; and key-controlled protection does not apply totracing, the second operand of Test Block, or instructions that operatespecifically on the linkage stack, whereas low address protection does.

Because fetch-protection override and low address protection do notapply to an address space for which the private-space control is one inthe address-space-control element, locations 0-2047 and 4096-4607 in theaddress space are usable the same as the other locations in the space.

Suppression-on-Protection

Some instruction definitions specify that the operation is suppressed ifa protection exception due to any type of protection is recognized. Whenthat specification is absent, the execution of an instruction issuppressed if a protection exception due to access-list-controlledprotection or DAT protection (a.k.a., page protection) is recognized,and it may be either suppressed or terminated if a protection exceptiondue to low-address protection or key-controlled protection isrecognized.

The suppression-on-protection function allows the control program tolocate the segment-table entry, page-table entry, and, when enhanced DATapplies, the region-table entry used in the translation of a virtualaddress that caused a protection exception, in order to determine if theexception was due to DAT protection. This is used, for instance, for theimplementation of the POSIX fork function (or copy-on-write function).The function also allows the control program to avoid locating thesegment-table and page table entries if the address was not virtual orthe exception was due to access-list-controlled protection. Whenenhanced DAT applies, the control program may also avoid locating theregion-table entries, if the address was not virtual or the exceptionwas due to access-list-controlled protection.

During a program interruption due to a protection exception, either aone or a zero is stored in, for instance, bit position 61 of reallocations 168-175. As one example, in the z/Architecture®, the storingof a one in bit position 61 indicates that:

-   -   The unit of operation or instruction execution during which the        exception was recognized was suppressed.    -   If dynamic address translation (DAT) was on, as indicated by the        DAT-mode bit in the program old PSW, the effective address that        caused the exception is one that was to be translated by DAT.        (The effective address is the address which exists before any        transformation by DAT or prefixing.) Bit 61 is set to zero if        DAT was on, but the effective address was not to be translated        by DAT because it is a real address. If DAT was off, the        protection exception cannot have been due to DAT protection or        access-list-controlled protection.    -   Bit positions 0-51 of real locations 168-175 contain bits 0-51        of the effective address that caused the exception. If DAT was        on, indicating that the effective address was to be translated        by DAT, bit positions 62 and 63 of real locations 168-175, and        real location 160, contain the same information as is stored        during a program interruption due to a page-translation        exception—this information identifies the address space        containing the protected address. Also, bit 60 of real locations        168-175 is zero if the protection exception was not due to        access-list-controlled protection or is one if the exception was        due to access-list controlled protection. A one in bit position        60 indicates that the exception was not due to DAT protection.        If DAT was off, the contents of bit positions 60, 62, and 63 of        real locations 168-175, and the contents of real location 160,        are unpredictable. The contents of bit positions 52-59 of real        locations 168-175 are unpredictable.

Bit 61 being zero indicates that the operation was either suppressed orterminated and that the contents of the remainder of real locations168-175, and of real location 160 are unpredictable.

Bit 61 is set to one if the protection exception was due toaccess-list-controlled protection or DAT protection. Bit 61 may be setto one if the protection exception was due to low-address protection orkey controlled protection.

If a protection-exception condition exists due to eitheraccess-list-controlled protection or DAT protection, but also exists dueto either low-address protection or key-controlled protection, it isunpredictable whether bit 61 is set to zero or one.

The suppression-on-protection function is useful in performing the POSIXfork function, which causes a duplicate address space to be created. Thefollowing discussion pertains to when enhanced DAT does not apply, orwhen enhanced DAT applies, but the format-control (FC) bit in thesegment-table entry is zero. When forking occurs, the control programcauses the same page of different address spaces to map to a single pageframe of real storage as long as a store in the page is not attempted.Then, when a store is attempted in a particular address space, thecontrol program assigns a unique page frame to the page in that addressspace and copies the contents of the page to the new page frame. Thislast action is sometimes called the copy-on-write function. The controlprogram sets the DAT-protection bit to one in the page-table entry for apage in order to detect an attempt to store in the page. The controlprogram may initially set the DAT-protection bit to one in asegment-table entry to detect an attempt to store anywhere in thespecified segment. When enhanced DAT applies, and the format control(FC) bit in the segment-table entry is one, a similar technique may beused to map a single segment frame of absolute storage.

Bit 61 being one in real locations 168-175 when DAT was on indicatesthat the address that caused a protection exception is virtual. Thisindication allows programmed forms of access register translation anddynamic address translation to be performed to determine whether theexception was due to DAT protection as opposed to low-address orkey-controlled protection.

The results of suppression on protection are summarized in the tablebelow:

Presented Fields Exception Conditions If Bit 61 One LA or Key- ALC orBits 62, 63 Cont. Prot. DAT Page Prot. Eff. Addr. Bit 61 and Loc. 160Bit 60 No On Yes Log. 1 P 1A Yes On Yes Log. U1 P 1A Yes Off No Log. U2U3 U3 Yes Off No Real U2 U3 U3 Yes On No Log. U2 P 0 Yes On No Real OR —— Explanation: — Immaterial or not applicable. OR Zero because effectiveaddress is real. 1A One if bit 61 is set to one because ofaccess-list-controlled protection; zero otherwise. ALCAccess-list-controlled. LA Low-address. Log. Logical. P Predictable. U1Unpredictable because low-address or key-controlled protection may berecognized instead of access-list-controlled or page protection. U2Unpredictable because bit 61 is only required to be set to one foraccess-list-controlled or page protection. U3 Unpredictable because DATis off.Enhanced Suppression on Protection

When the enhanced suppression-on-protection function is installed, thereare the following additional constraints on what may occur during aprotection exception. These constraints take precedence over anyconstraints defined in the original suppression-on-protection function.

During a program interruption due to a protection exception, bit 61 ofreal locations 168-175 indicates the type of protection exceptionrecognized. Bit 61 is set to one if the protection exception was due toaccess-list-controlled protection or DAT protection. Bit 61 is set tozero if the protection exception was due to low-address protection orkey-controlled protection.

If a protection-exception condition exists due to eitheraccess-list-controlled protection or DAT protection, but also exists dueto either low-address protection or key-controlled protection, it isunpredictable which exception is recognized and whether bit 61 is set tozero or one. However, while it is unpredictable which exception isrecognized, the recognized exception produces consistent behavior, assummarized in the below table:

Exception Bit Bit Bits 62, 63 Bits Type DAT 61 60 and Loc. 160 0-51 LAPAny 0 — — — KCP Any 0 — — — ALCP On 1 1 AS A DATP On 1 0 AS AExplanation: A Bits 0-51 of the effective address that caused theexception. ALCP Access-list-controlled protection. AS Identifies theaddress space containing the effective address that caused theexception. DATP DAT protection. KCP Key-controlled protection. LAPLow-address protection. — Undefined.

The SOP and ESOP facilities are provided for a native architecture, aswell as in interpretative execution mode.

In the z/Architecture®, suppression-on-protection in theinterpretive-execution mode is used to handle protection exceptions whenan exception control (e.g., ECA.18) is one; and enhancedsuppression-on-protection in the interpretive-execution mode is used forbehavior when an exception control (e.g., ECA.18) is one and theenhanced suppression-on-protection facility is installed. In thisexample, ECA.18 is the protection-interception control, and in general,if a bit in ECA is zero, the associated condition results ininterception, and if a bit in ECA is one, interpretive execution of theassociated function is attempted by the machine.

Further, in one example, an execution control (e.g., ECB.6) is thehost-protection-interruption control. In general, if a bit in ECB iszero, the associated condition results in interception, and if a bit inECB is one, interpretive execution of the associated function isattempted. Enhanced suppression-on-protection in theinterpretive-execution mode is used for the handling of host pageprotection exceptions when ECB.6 is one and the enhancedsuppression-on-protection facility is installed.

A protection exception may result in a program interruption interceptiondepending on the values of ECA.18, ECB.6, and IC.2 (an interceptioncontrol). Other program interruptions cause interception with, forinstance, an interception code 08, if they are of a type for which thecorresponding interception control, bit 1 or bit 2, is set to one. Withinterception code 08, interception is mandatory, in one example, forguest program interruptions caused by, for instance, these exceptions:addressing, specification, and special operation.

A guest addressing exception may be due to an invalid guest or hostaddress.

A guest protection exception may be due to a guest protection condition,a host page-protection condition, or in a given mode (e.g.,MCDS—Multiple Control Data Spaces), a host access-list controlledprotection condition.

The parameters of the interruption are placed in the state description.

Guest accesses to guest storage, in the various modes of interpretiveexecution, are subject to key-controlled storage protection as definednatively, using the real-machine storage keys. Guest accesses are alsosubject to guest low-address protection, guest page protection, guestfetch-protection override, and guest access-list-controlled protection.Host low-address protection is not applied to guest references to gueststorage.

In addition, host page protection applies to pageable-storage-mode-guestreferences, as does host access-list-controlled protection in the MCDSmode.

In general, disallowed storing causes a protection exception to berecognized, which results in a guest or host program interruption.However, store access to the first 4K-byte block of the guest prefixarea is checked on entry to the interpretive-execution mode, and anyaccess exception encountered during this test results in validityinterception (e.g., VIR code 0037 hex). Subsequently, duringinterpretive execution, an access exception condition encountered onaccess to the first 4K-byte block of the guest prefix area may bepresented normally, or may result in a validity interception. If anaccess exception condition encountered by a reference to the guestprefix area to perform a guest interruption results in validityinterception, then the following information may be lost:

-   -   Interruption information to be stored in the guest prefix area.        (However, the guest PSW prior to the interruption is stored in        the state description.)    -   The interruption request itself, if applicable.

The host page-protection and access-list-controlled-protectionmechanisms are taken into consideration when the condition code is setduring the interpretive execution of a guest Test Protectioninstruction, described below.

Further details regarding SOP and ESOP in interpretative execution modeare described below.

Suppression-on-Protection in Interpretive-Execution Mode

The suppression-on-protection function is usable by, for instance,z/VM®, by means of host access-list-controlled protection and host pageprotection, for a pageable-mode DAT-off guest. A guest is DAT-off whenthe DAT-mode bit (e.g., bit 5 of the program-old PSW) is zero; a guestis DAT-on when the DAT mode bit is one. Host access-list-controlledprotection applies to an MCDS guest (which is a DAT-off guest).

In the interpretive-execution mode, for the purpose of thissuppression-on-protection definition, the effective address stored inreal locations 168-175 is the guest effective address, even when theprotection exception is due to host access-list-controlled protection orhost page protection.

When a protection exception is caused by host access-list-controlledprotection or host page protection and the guest is a DAT-off guest,then, even when the guest effective address is one that is defined to bea real (as opposed to logical) address, bit 61 is set to one, and bits0-51 of the guest effective address are stored. In this case, bits 62and 63 of real locations 168-175, and the contents of real location 160,properly indicate the host address space to which the exception appliesif the guest is an MCDS-mode guest, or the contents of those fields areunpredictable if the guest is a non-MCDS-mode DAT-off guest. (In thelatter case, it is known that the exception applies to the host primaryaddress space.)

Without ESOP, when a protection exception is caused by host pageprotection and the guest is a DAT-on guest, then bit 61 isunpredictable. Host page protection is not usable with DAT-on guestsbecause of the inability to distinguish between a DAT-on guestpage-protection exception versus a host page-protection exception.

As is normal, the above references to real locations refer to guest reallocations if a guest interruption occurs. If an interception occursinstead of a guest interruption, information is stored instead atcorresponding locations in the state description.

As described above, ECA.18 (in the state description) is the protectioninterception control. When ECA.18 is zero, a protection exceptionresults in mandatory interception. When ECA.18 is one, interception isprevented and a guest interruption occurs. ECA.18 is overridden by IC.2,which specifies, when one, that interception occurs on any programinterruption. ECA.18 is retainable by the SIE state-retention assist.

Enhanced Suppression-on-Protection in Interpretive-Execution Mode

Enhanced suppression-on-protection in interpretive-execution mode isdefined, for instance, for hosts running in the z/Architecture® mode andguests running in any architectural mode under hosts in thez/Architecture® mode. (In other embodiments, however, the hosts can berunning in other architectural modes.)

ECA.18 and ECB.6 (in the state description) are theprotection-interception and host-protection interruption controls,respectively. When ECA.18 is one and IC.2 is zero, interception isprevented for any protection exception recognized by the guest, and aguest interruption occurs. When ECA.18 is zero, protection exceptions(except host page protection exceptions when ECB.6 is one) result inmandatory interception.

If ECB.6 is one, exceptions due to host page-protection are presented ashost program interruptions. If ECB.6 is zero, host page-protectionexceptions are presented as guest program interruptions orinterceptions, as governed by ECA.18 and IC.2. In the case of a hostinterruption due to host page protection, the following occurs:

-   -   The guest unit of operation or guest instruction execution        during which the exception was recognized is nullified.    -   The Start Interpretive Execution instruction is suppressed.    -   Information is stored in host storage as though it were a normal        exception in the host. The address stored in the        translation-exception ID is the host virtual address, and the        ASCE ID and exception access ID identify the host address space.        (For a non-MCDS guest, this is the host primary space.)

Interception prevention due to ECA.18 is overridden by IC.2, whichspecifies, when one, that interception occurs, instead of a guestprogram interruption. However, if both ECB.6 and IC.2 are one, hostpage-protection exceptions are still presented as host interruptions(that is, the ECB.6=1 treatment takes precedence over the IC.2=1treatment). ECA.18 and ECB.6 are retainable by the SIE state-retentionassist.

When a guest protection exception is caused by hostaccess-list-controlled protection or host page protection and the guestis in the MCDS mode, then, even when the guest effective address isdefined to be a real (as opposed to a logical) address, bit 61 of guestreal locations 168-175 is set to one, bits 0-51 of the guest effectiveaddress are stored, and bits 62 and 63 and the contents of real location160 properly indicate the host address space to which the exceptionapplies. When a protection exception due to host page protection isrecognized in a non-MCDS guest and host-protection interruption isdisabled (that is, ECB.6 is zero), then bit 61 of guest real locations168-175 is set to zero, and the remaining fields are undefined.

Further details regarding SOP and ESOP are described with reference toFIGS. 8A-8B, in which one embodiment of SOP/ESOP guest processing isdescribed, in accordance with an aspect of the present invention.

Referring initially to FIG. 8A, entry to this logic is based on a guestinstruction, issued by a pageable guest executing within theenvironment, that attempts to perform a fetch or store access at astorage location, STEP 800. In response to a request for storage access,a determination is made as to whether the requested access (e.g., fetchor store) is permitted by the guest and the host, INQUIRY 802. In thisexample, the access is attempted, and if successful, the access isperformed and the instruction completes, STEP 804. However, if theattempted access fails, then processing continues to determine whetherit was host level protection or guest level protection that caused theaccess to fail.

If the access is not permitted, a determination is made as to whetherDAT processing is on for the guest (a.k.a., DAT-on guest), INQUIRY 806.If the guest is not operating under translation (i.e., it is a DAT-offguest), then a further inquiry is made as to whether key controlledprotection is set for the guest, INQUIRY 808. In this example, for aDAT-off guest, there still may be guest key controlled protection of areal storage block. If there is no guest key controlled protection, thenit is not a guest-level protection denying access, but instead, a hostlevel protection. Thus, a host interrupt is provided, as describedbelow, STEP 810.

Returning to INQUIRY 808, if there is guest key controlled protection,then a guest interrupt is presented, STEP 812. This indicates to theguest that it tried to access a location that the guest indicated wasprotected.

Referring again to INQUIRY 806, if DAT is on for the guest, then afurther inquiry is made as to whether there are other protection formsthat may also apply including, for instance, low address protection,access list controlled protection, key controlled protection or DATprotection (also known as page protection prior to enhanced DAT). Ifnone of these other protection forms apply, then it is a host problem(i.e., a host level protection causing access to fail), and thus a hostinterrupt is presented, STEP 810. However, if one of these otherprotection forms apply, then a guest interrupt is provided, STEP 816.

Further, if a guest protection scheme applies, a determination is madeas to whether it is guest DAT protection (DATP) that applies and whetherthe request is for a store access, INQUIRY 818. If not, then a guestexception is processed, STEP 820. In one example, the supervisor (e.g.,host) performs whatever processing is needed or desired, as determinedby policy, when a guest receives one of these protection conditions.

Returning to INQUIRY 818, if it is a guest DATP and a store access, thenpossibly, the guest operating system understands copy-on-write at theguest level, STEP 822. If there is no copy-on-write notion, then a guestexception may be performed.

Further details regarding handling the host interrupt are described withreference to FIG. 8B. Initially, a determination is made as to whetherthe protection interrupt is due to a host DATP on a store access,INQUIRY 840. If not, then the condition is reflected to the guest, STEP842. In one example, the host performs whatever actions are necessary toindicate to the guest the particular condition causing the fault.

Referring again to INQUIRY 840, if it is a host DATP on a store access,then a determination is made as to whether it is a DAT-on guest, INQUIRY844. Should it be a DAT-on guest, then a further inquiry is made as towhether SOP or ESOP processing is controlling, INQUIRY 846. If it is SOPprocessing, then the condition is reflected to the guest, STEP 848.However, if it is ESOP processing, then bit 61 in the translationexception identification (TEID) is predictable for the store accessattempt that encountered the DATP condition. Without ESOP, guest DAT-onleads to the reflect to guest conclusion.

If the guest DAT is off, INQUIRY 844, or if the guest DAT is on and itis ESOP, INQUIRY 846, then the following occurs. A determination is madeas to whether the storage area to be accessed (e.g., page) is part ofthe host copy-on-write scheme, INQUIRY 850. If not, then the conditionis reflected to the guest, STEP 852. However, if it is part of the hostcopy-on-write scheme, then the host copy-on-write is performed, STEP854.

Advantageously, ESOP provides the means by which a host and a guest knowwhat is going on for a storage (e.g., page) fault and who gets blamed.The ESOP facility provides the ability of the machine (e.g., CPU) todiscern guest from host page protection when operating interpretively.Previously, page protection could not be distinguished from being blamedon guest vs. host page fault. That is, a page-protection exception couldbe either guest or host and some unpredictability existed. Thus, thehost could not do much with it. With ESOP, the prior unpredictability iseliminated and a host page-protection causes a host program interruptionwith information about the host bad page (e.g., failing host address).Similarly, a guest page-protection produces a guest program interruptionwith information about the guest page (e.g., failing guest address).

Previously, a host COW could be performed only for a guest that the hostprogram knew was a DAT-off guest (e.g. CMS, MCDS) because for a DAT-offguest, the page protection must have been a host matter. For a DAT-onguest, however, it was unclear whether the page protection was a host orguest matter, and therefore, COW could not be used.

In accordance with an aspect of the present invention, this has changed.The enhanced suppression-on-protection (ESOP) facility enables adistinction to be made between host-level protection and guest levelprotection, and therefore, enables the use of COW, when appropriate.With ESOP, the additional function allows the distinction to be made sothat even under circumstances of interception for a DAT-On guest (thatmight have tripped over its own protected page), sufficient informationis available to tell the host whose exception it is, and to stillperform copy-on-write, when it does belong to the host.

Test Protection

In accordance with an aspect of the present invention, a capability isprovided that enables testing for protection exceptions. As one example,a Test Protection (TPROT) instruction is used to provide to a programexecuting the TPROT instruction information about storage to be used bythe program. For instance, it provides indications about the protectionsof the storage and allows distinctions to be made between host levelprotections and guest level protections. TPROT takes advantage of ESOP,in accordance with one or more aspects of the present invention.

To further explain, when a program allocates a storage frame,frequently, it will execute a query command to learn if the frame hasattributes allowing access by the program (i.e., is it resident, is itnot write-protected). As examples, in IBM® System Z® processors, theTEST Protection instruction is used for this purpose, and in SunMicrosystems, a mincore( ) system call is used. By issuing theinstruction or call, the program learns the attributes of the framethrough a condition code in the former, and a return code in the latter.When the program issuing the storage query command is executinginterpretively as a guest under the control of a host hypervisor, thereturn code may reflect the attributes of the frame from the viewpointof the guest address translation tables, as well as the host addresstranslation tables. Such behavior may expose the effects of using aparticular host storage control scheme, such as copy-on-write to theguest, thus compromising the isolation provided by a hypervisor.

As described herein, the copy-on-write (COW) scheme is used to reducememory pressure by operating systems. As processor performanceincreases, so too does the ability to run multiple programs under thecontrol of a single operating system, and multiple operating systemsunder the control of a single hypervisor. To accommodate a larger numberof simultaneous programs, special measures are implemented to managesystem memory requirements.

One of these measures is the use of COW. When a program is started bythe operating system, its initially allocated storage frames are markedas write-protected, and are physically shared with other programs underthe control of the same OS. Then, when a store operation is attempted toone of these frames, a program exception occurs and the operating systemallocates a unique frame of storage and copies in the contents of theoriginal shared frame.

Some programs may operate successfully with a storage query command thatreturns the combined view of a frame's attributes at the guest and hostlevel. However, if a program is written to avoid allocatingwrite-protected storage, it may fail if the storage query commandreturns an indication of write-protection for all queried frames, havingno recourse through modification of the guest's view of the frame. Sucha case may occur if the hypervisor does not implement a copy-on-writescheme.

Some systems may allow guests programs to be written to usepara-virtualization; guest requirements are passed to the host through adefined interface. In this case, the guest may specify directly to thehost its intention of using storage and the host may respond asdescribed above by allocating a copied frame. In such a case, the guestprogram is to be modified to account for the interface to eachhypervisor host that it executes on. Also, the para-virtualization callstypically have a high overhead, and cause execution pressure on thehypervisor.

When an interpretively executing guest program issues a query commandagainst a storage frame, if conditions exist such that the conditioncode or return code of the command would indicate the frame iswrite-protected solely because it is write-protected in the host addresstranslation tables, then a program interception occurs, instead ofexecution completing with a condition code or return code indicationthat reflects the state of the host address translation table.

In accordance with an aspect of the present invention, certain hoststorage management techniques, such as copy-on-write, operatetransparently to a guest program. It allows program logic to rely on theresult of a storage query command to reflect only the guest view of aframe's storage attributes. In the case that a page is write-protectedonly in the guest address translation tables, the guest will learn of itwithout having to go through a system call or through communication withthe hypervisor.

In one example, the Test Protection (TPROT) instruction takes a frameaddress as an input operand, and completes with one of the followingfour return codes:

-   -   0 Fetching permitted; storing permitted    -   1 Fetching permitted; storing not permitted    -   2 Fetching not permitted; storing not permitted    -   3 Translation not available.

In accordance with an aspect of the present invention, a facility isprovided for executing a Test Protection (TPROT) instruction such thatcertain benefits are provided including, but not limited to,distinguishing between host level and guest level protection; andproviding a false return code that indicates execution completedsuccessfully, although the translation tables are not updated toaccommodate COW. In this embodiment, a hypervisor sets a state value(ECB.6) for a guest image of a logical partition; a guest program of theguest image fetches the TPROT instruction for execution, the TPROTinstruction specifying a memory location to be tested; the TPROTinstruction tests the memory location for protection (store); responsiveto the state value indicating no-suppression, performing any of settinga condition code based on a test, or performing an exception operation;or responsive to the state value indicating suppression, performingresponsive to the test indicating a condition code should be set,suppressing execution of the instruction; and performing an interceptionof the instruction.

The execution of a Test Protection instruction by a pageable-mode guestis suppressed and a host interruption is recognized (instead of settingcondition code 1) when all of the following are true, in one example:

-   -   The enhanced-suppression-on-protection facility is installed in        the host.    -   ECB.6 is one.    -   Conditions exist which call for condition code 1 to be set.    -   A host page-protection-exception condition exists.    -   Neither an access-list-controlled-protection exception nor a        guest page-protection exception condition exists.

During a storage reference, when a guest page protection condition isencountered, a condition code of 1 is set by the guest unless conditioncode 2 is set because of key-controlled storage protection, and noaction is taken by the host. For a host page protection, execution bythe guest is suppressed and an instruction interception is recognized,if the conditions specified above are true, and no action is taken bythe host.

One example of a format of the TPROT instruction is described withreference to FIG. 9. As depicted, a TPROT instruction 900 includes, forinstance, an opcode 902 (e.g., E501) identifying the TPROT instruction;a first register field (B₁) 904 and a first displacement field (D₁) 906;and a second register field (B₂) 908 and a second displacement field(D₂) 910. A first operand address is determined by adding the contentsof the D₁ field to the contents of a register identified by the B₁field. Likewise, a second operand address is formed by adding thecontents of the D₂ field with the contents of a register specified bythe B₂ field.

In execution, the location designated by the first-operand address istested for protection exceptions by using the access key specified by,for instance, bits 56-59 of the second operand address.

The second-operand address is not used to address data; instead, bits56-59 of the address form the access key to be used in testing. Bits0-55 and 60-63 of the second-operand address are ignored, in thisexample.

The first-operand address is a logical address. When the CPU is in theaccess-register mode (when DAT is on and PSW bits 16 and 17 are 01binary), the first operand address is subject to translation by means ofboth the access-register-translation (ART) and thedynamic-address-translation (DAT) processes. ART applies to the accessregister designated by the B1 field, and it obtains theaddress-space-control element to be used by DAT. When DAT is on, but theCPU is not in the access-register mode, the first operand address issubject to translation by DAT. In this case, DAT uses theaddress-space-control element contained in control register 1, 7, or 13when the CPU is in the primary-space, secondary-space, or home-spacemode, respectively. When DAT is off, the first-operand address is a realaddress not subject to translation by either ART or DAT.

When the CPU is in the access-register mode and an address-space-controlelement cannot be obtained by ART because of a condition that wouldnormally cause one of the exceptions shown in the following table, theinstruction is completed by setting condition code 3.

Exception Name Cause ALET specification Access-list-entry-token (ALET)bits 0-6 not all zeros ALEN translation Access-list entry (ALE) outsidelist or invalid (bit 0 is one) ALE sequence ALE sequence number (ALESN)in ALET not equal to ALESN in ALE ASTE validity ASN-second-table entry(ASTE) invalid (bit 0 is one) ASTE sequence ASTE sequence number(ASTESN) in ALE not equal to ASTESN in ASTE Extended authority ALEprivate bit not zero, ALE authorization index (ALEAX) not equal toextended authorization index (EAX), and secondary bit selected by EAXeither outside authority table or zero

When the access register contains 00000000 hex or 00000001 hex, ARTobtains the address-space-control element from control register 1 or 7,respectively, without accessing the access list. When the B₁ fielddesignates access register 0, ART treats the access register ascontaining 00000000 hex and does not examine the actual contents of theaccess register.

When ART is completed successfully, the operation is continued throughthe performance of DAT.

When DAT is on and the first-operand address cannot be translatedbecause of a condition that would normally cause one of the exceptionsshown in the following table, the instruction is completed by settingcondition code 3.

Exception Name Cause ASCE type Address-space-control element (ASCE)being used is a region-second-table designation, and bits 0-10 offirst-operand address not all zeros; ASCE is a region-third tabledesignation, and bits 0-21 of first-operand address not all zeros; orASCE is a segment-table designation, and bits 0-32 of first-operandaddress not all zeros. Region first translation Region-first-table entryoutside table or invalid. Region second Region-second-table entryoutside table or translation invalid. Region third translationRegion-third-table entry outside table or invalid. Segment translationSegment-table entry outside table or invalid Page translation Page-tableentry invalid

When translation of the first-operand address can be completed, or whenDAT is off, the storage key for the block designated by thefirst-operand address is tested against the access key specified in bitpositions 56-59 of the second-operand address, and the condition code isset to indicate whether store and fetch accesses are permitted, takinginto consideration the applicable protection mechanisms. Thus, forexample, if low-address protection is active and the first-operandeffective address is in the range 0-511 or 4096-4607, then a storeaccess is not permitted. Access-list-controlled protection, pageprotection, storage-protection override, and fetch-protection overridealso are taken into account.

The contents of storage, including the change bit, are not affected.Depending on the model, the reference bit for the first-operand addressmay be set to one, even for the case in which the location is protectedagainst fetching.

When the CPU is in the access-register mode, an addressing exception isrecognized when the address used by ART to fetch the effective accesslist designation or the ALE, ASTE, or authority-table entry designates alocation which is not available in the configuration.

When DAT is on, an addressing exception is recognized when the addressof the region-table entry or entries, segment-table entry, or page-tableentry or the operand real address after translation designates alocation which is not available in the configuration. Also, atranslation-specification exception is recognized when a region-tableentry or the segment-table entry or page-table entry has a format error.When DAT is off, only the addressing exception due to the operand realaddress applies, in this example.

For all of the above cases, the operation is suppressed.

Resulting Condition Code:

0 Fetching permitted; storing permitted

1 Fetching permitted; storing not permitted

2 Fetching not permitted; storing not permitted

3 Translation not available

Program Exceptions:

-   -   Addressing (effective access-list designation, access-list        entry, ASN-second-table entry, authority-table entry,        region-table entry, segment-table entry, page-table entry, or        operand 1).    -   Privileged operation.    -   Translation specification.

Test Protection permits a program to determine the protection attributesof an address passed from a calling program without incurring programexceptions. The instruction sets a condition code to indicate whetherfetching or storing is permitted at the location designated by thefirst-operand address of the instruction. The instruction takes intoconsideration the protection mechanisms in the machine: e.g.,access-list controlled, page, key-controlled, low address protection,storage-protection override, and fetch-protection override.Additionally, since ASCE-type, region-translation, segment-translation,and page-translation-exception conditions may be a program substitutefor a protection violation, these conditions are used to set thecondition code rather than cause a program exception. When the CPU is inthe access-register mode, Test Protection additionally permits theprogram to check the usability of an access-list entry token (ALET) inan access register without incurring program exceptions. The ALET ischecked for validity (absence of an ALET-specification,ALEN-translation, and ALE-sequence exception condition) and for beingauthorized for use by the program (absence of an ASTE-validity,ASTE-sequence, and extended-authority exception condition).

The approach using Test Protection has the advantage of a test whichdoes not result in interruptions; however, the test and use areseparated in time and may not be accurate if the possibility exists thatthe storage key of the location in question can change between the timeit is tested and the time it is used.

In the handling of dynamic address translation, Test Protection issimilar to Load Real Address in that the instructions do not causeASCE-type, region-translation, segment-translation, and page-translationexceptions. Instead, these exception conditions are indicated by meansof a condition-code setting. Similarly, access-register translation setsa condition code for certain exception conditions when performed duringeither of the two instructions. Conditions which result in conditioncodes 1, 2, and 3 for Load Real Address, result in condition code 3 forTest Protection. The instructions also differ in several other respects.The first-operand address of Test Protection is a logical address andthus is not subject to dynamic address translation when DAT is off. Thesecond operand address of Load Real Address is a virtual address whichis translated.

Access-register translation applies to Test Protection when the CPU isin the access-register mode (DAT is on), whereas it applies to Load RealAddress when PSW bits 16 and 17 are 01 binary regardless of whether DATis on or off. When condition code 3 is set because of an exceptioncondition in access-register translation, Load Real Address, but notTest Protection, returns in a general register the program-interruptioncode assigned to the exception.

Condition code 3 does not necessarily indicate that the first-operandlocation will always be inaccessible to the program; rather it merelyindicates that the current conditions prevent the instruction fromdetermining the protection attributes of the operand. For example, in avirtual storage environment, condition code 3 may be set if the storagelocation has been paged out by the operating system. If the programattempts to access the location, the operating system may resolve thepage-translation exception and subsequently make the location accessibleto the program. Similarly, condition code 1 does not necessarilyindicate that the address cannot ever be stored into. In an operatingsystem that implements a Posix fork function, page protection is used toalert the operating system of a copy-on-write event. Following theoperating-system resolution of the copy-on-write event, the program maybe given store access to the location.

One embodiment of the logic associated with executing a Test Protectioninstruction, in accordance with an aspect of the present invention, isdescribed with reference to FIGS. 10A-10B. Initially, a guest executesthe Test Protection instruction in order to learn information about thestorage being queried by the Test Protection instruction, STEP 1000.During execution, a determination is made as to whether a store ispermitted by the guest and the host, INQUIRY 1002. In one example, thisdetermination is made by testing whether the store could occur, althoughthe store does not actually occur. For instance, various permissions andtests are checked with the address of the store, but the store is notperformed. If the store would be successful, if performed, then thestoring is considered permitted by both the guest and the host. Thus,the query instruction completes with a successful condition code (e.g.,zero), STEP 1004.

However, if storing is not permitted by the guest, the host or both,then a further inquiry is made as to whether guest address translationis available, INQUIRY 1006. As examples, if the guest is a DAT-offguest, then translation is available in the sense that the identity isthe translated address. However, if the guest is a DAT-on guest, one ormore guest translation tables are searched to see if the correspondingentry is valid. If it is valid, then the guest address translation isavailable for the DAT-on guest. If guest address translation isunavailable, then the Test Protection instruction completes with acondition code indicating this situation (e.g., CC=3), STEP 1008.

If guest address translation is available, a further determination ismade as to whether host address translation is available, INQUIRY 1010.This determination is made similarly to determining whether the guestaddress translation is available. However, for a DAT-on host, the DATtranslation tables are used, instead of the guest translation tables. Ifhost address translation is unavailable, then a host interrupt isexecuted, STEP 1012. In one example, the guest exits and there is a hostexception routine that is invoked in order to manage the situation.

If both guest address translation and host address translation areavailable, then a further determination is made as to whether fetchingis permitted by the guest and host, INQUIRY 1014. In one example, thisdetermination is made based on the permissions stored at the DAT tables.If fetching is not permitted by both the guest and the host, then theinstruction completes with a selected condition code (e.g., CC=2), STEP1016. If fetching is permitted by both the guest and the host, then afurther determination is made as to whether the host protectionexception condition trap is enabled, INQUIRY 1018. In one example, thehost protection exception condition, which is an aspect of the presentinvention, is enabled when ESOP is enabled and ECB.6 is on. If the hostprotection exception condition trap is not enabled, then the instructioncompletes with a selected condition code (e.g., CC=1), STEP 1020.Without the trap enabled, no special distinction is made between guestand host, nor between DAT and other types of store protection.

If the host protection exception condition trap is enabled, then afurther determination is made as to whether storing is permitted by theguest settings, INQUIRY 1022. In one example, this determination is madeby checking the settings set by the guest in, for instance, translationtables (e.g., guest DAT table) and/or other data structures such asthose that include the protections set by the guest. If storing is notpermitted, then the instruction completes with a selected condition code(e.g., CC=1), STEP 1024. However, if storing is permitted by the guest,then a further determination is made as to whether the host writeprotection is other than DAT, INQUIRY 1026. If so, then the instructioncompletes with a specified condition code (e.g., CC=1), STEP 1028.However, if only host DAT protection exists, then a host interception isprovided, STEP 1030.

At this point, the TPROT instruction is simulated by the host, STEP1032. That is, when guest execution of an instruction causes aninterception to the host, the host program modifies the guestenvironment so that when the guest program is restarted (by the host),it appears to the guest that the intercepted instruction has executed.During the simulation, a determination is made as to whether the area instorage to be accessed (e.g., page) is part of the host COW scheme,INQUIRY 1040. If the page is not part of the host COW scheme indicatingthat it cannot be written to, then a selected completion code isprovided (e.g., CC=1), STEP 1042. However, if the page is part of thehost COW scheme, then TPROT issues with the successful completion code(e.g., CC=0), STEP 1044.

In accordance with an aspect of the present invention, with thissuccessful condition code, the host does not change the state of its DATtables. The frame remains write-protected. In a host COW scheme, whenthe guest actually stores to the frame, a host interruption will occur(e.g., via ESOP) and the host exception routine will perform the COW.This “false” return code is an advantage of one aspect of the inventionin that when TPROT is issued against a block that is never stored to,performing the COW is avoided.

In one embodiment, a protection query instruction (e.g., TPROT) isprovided, in which when the protection setting in the host table isfetch protected, a host interrupt occurs. However, when the protectionsetting in the guest table is fetch protected, a condition code isreturned.

As one example, if the TPROT instruction is interpretively executed, andif it would complete with CC=1 only because the frame is write-protectedwithin the host view (i.e., in the host address translation table) andnot in the guest view (i.e., in the guest address translation table),then a host interception occurs. While handling the interception, if thehost is implementing a copy-on-write (COW) storage control scheme and ifthe frame is truly available to be modified by a guest program, a copywill be made of the frame, the host address translation tables will beupdated, and guest execution will continue with the TPROT instructioncompleting with CC=0.

One or more aspects of the present invention can be included in anarticle of manufacture (e.g., one or more computer program products)having, for instance, computer usable media. The media has therein, forinstance, computer readable program code means or logic (e.g.,instructions, code, commands, etc.) to provide and facilitate thecapabilities of the present invention. The article of manufacture can beincluded as a part of a computer system or sold separately.

One example of an article of manufacture or a computer program productincorporating one or more aspects of the present invention is describedwith reference to FIG. 11. A computer program product 1100 includes, forinstance, one or more computer usable media 1102 to store computerreadable program code means or logic 1104 thereon to provide andfacilitate one or more aspects of the present invention. The medium canbe an electronic, magnetic, optical, electromagnetic, infrared, orsemiconductor system (or apparatus or device) or a propagation medium.Examples of a computer readable medium include a semiconductor or solidstate memory, magnetic tape, a removable computer diskette, a randomaccess memory (RAM), a read-only memory (ROM), a rigid magnetic disk andan optical disk. Examples of optical disks include compact disk-readonly memory (CD-ROM), compact disk-read/write (CD-R/W) and DVD.

A sequence of program instructions or a logical assembly of one or moreinterrelated modules defined by one or more computer readable programcode means or logic direct the performance of one or more aspects of thepresent invention.

Commercial Implementation

Although the z/Architecture® by IBM® is mentioned herein, one or moreaspects of the present invention are equally applicable to other machinearchitectures and/or computing environments employing pageable entitiesor similar constructs.

Commercial implementations of the TPROT instruction, facilities, andother formats, instructions, and attributes disclosed herein can beimplemented either in hardware or by programmers, such as operatingsystem programmers, writing in, for example, assembly language. Suchprogramming instructions may be stored on a storage medium intended tobe executed natively in a computing environment, such as az/Architecture® IBM® server, or alternatively in machines executingother architectures. The instructions can be emulated in existing and infuture IBM® servers and on other machines or mainframes. They can beexecuted in machines where generally execution is in an emulation mode.

In emulation mode, the specific instruction being emulated is decoded,and a subroutine is built to implement the individual instruction, as ina subroutine or driver, or some other technique is used for providing adriver for the specific hardware, as is within the skill of those in theart after understanding the description hereof. Various software andhardware emulation techniques are described in numerous United Statespatents including: U.S. Pat. Nos. 5,551,013, 5,574,873, 5,790,825,6,009,261, 6,308,255, and 6,463,582, each of which is herebyincorporated herein by reference in its entirety. Many other teachingsfurther illustrate a variety of ways to achieve emulation of aninstruction format architected for a target machine.

In addition to the above, further details regarding guest processors andrelated processing is described in U.S. Pat. No. 7,197,585 entitled“Method and Apparatus for Managing the Execution of a BroadcastInstruction on a Guest Processor,” Farrell et al., issued Mar. 27, 2007,which is hereby incorporated herein by reference in its entirety.

Other Variations and Architectures

While various examples and embodiments are described herein, these areonly examples, and many variations are included within the scope of thepresent invention. For example, the computing environment describedherein is only one example. Many other environments, including othertypes of communications environments, may include one or more aspects ofthe present invention. For instance, different types of processors,guests and/or hosts may be employed. Further, pageable hosts, as well aspageable guests, may use one or more aspects of the present invention.Moreover, other types of architectures can employ one or more aspects ofthe present invention.

Aspects of the invention are beneficial to many types of environments,including environments that have a plurality of zones, andnon-partitioned environments. Further, there may be no central processorcomplexes, but yet, multiple processors coupled together. Variousaspects hereof are applicable to single processor environments.

Further, in the examples of the data structures and flows providedherein, the creation and/or use of different fields may include manyvariations, such as a different number of bits; bits in a differentorder; more, less or different bits than described herein; more, less ordifferent fields; fields in a differing order; different sizes offields; etc. Again, these fields were only provided as an example, andmany variations may be included. Further, indicators and/or controlsdescribed herein may be of many different forms. For instance, they maybe represented in a manner other than by bits. Additionally, althoughthe term address is used herein, any designation may be used.

As used herein, the term “page” is used to refer to a fixed-size orpredefined-size area of virtual storage (i.e., virtual memory). As oneexample, a host page is an area of host virtual storage. The size of thepage can vary, although in the examples provided herein, a page is 4Kbytes. Further, a “frame” is used to refer to a fixed-size or predefinedsize area of real or absolute storage (i.e., memory). As examples, ahost frame is an area of host real or absolute storage, and a guestframe is an area of guest real or absolute storage. In the case of apageable guest, this guest real or absolute storage is mapped by hostvirtual storage. As is common, pages of host virtual storage are backedby frames of host real or absolute storage, as needed. The size of theframe can vary, although in the examples provided herein, a frame is4K-bytes or 1M-bytes. However, in other embodiments, there may bedifferent sizes of pages, frames, segments, regions, blocks of storage,etc. Moreover, in other architectures, the terms “page” and “segment”may be used interchangeably or the term “page” may be used to apply tomultiple size units of virtual storage. The term “obtaining”, such asobtaining an instruction, includes, but is not limited to, fetching,having, receiving, being provided, creating, forming, issuing, etc. Aninstruction can reference other registers or can reference other thanregisters, such as operands, fields, locations, etc. Many otheralternatives to the above are possible. Further, although terms, such aslists, tables, etc. are used herein, any types of data structures may beused. For instance, a table can include other data structures as well.Again, those mentioned herein are just examples.

Further, a data processing system suitable for storing and/or executingprogram code is usable that includes at least one processor coupleddirectly or indirectly to memory elements through a system bus. Thememory elements include, for instance, local memory employed duringactual execution of the program code, bulk storage, and cache memorywhich provide temporary storage of at least some program code in orderto reduce the number of times code must be retrieved from bulk storageduring execution.

Input/Output or I/O devices (including, but not limited to, keyboards,displays, pointing devices, DASD, tape, CDs, DVDs, thumb drives andother memory media, etc.) can be coupled to the system either directlyor through intervening I/O controllers. Network adapters may also becoupled to the system to enable the data processing system to becomecoupled to other data processing systems or remote printers or storagedevices through intervening private or public networks. Modems, cablemodems, and Ethernet cards are just a few of the available types ofnetwork adapters.

The capabilities of one or more aspects of the present invention can beimplemented in software, firmware, hardware, or some combinationthereof. At least one program storage device readable by a machineembodying at least one program of instructions executable by the machineto perform the capabilities of the present invention can be provided.

The flow diagrams depicted herein are just examples. There may be manyvariations to these diagrams or the steps (or operations) describedtherein without departing from the spirit of the invention. Forinstance, the steps may be performed in a differing order, or steps maybe added, deleted, or modified. All of these variations are considered apart of the claimed invention.

Although embodiments have been depicted and described in detail herein,it will be apparent to those skilled in the relevant art that variousmodifications, additions, substitutions and the like can be made withoutdeparting from the spirit of the invention and these are thereforeconsidered to be within the scope of the invention as defined in thefollowing claims.

What is claimed is:
 1. A computer program product for facilitatingmanagement of storage of a computing environment, the computer programproduct comprising: a non-transitory computer readable storage mediumreadable by a processing circuit and storing instructions for executionby the processing circuit for performing a method comprising: attemptingaccess to a storage location, the storage location being protected by atleast one type of storage protection of one or more types of storageprotection, the one or more types of storage protection including atleast one of key controlled protection, access list controlledprotection, dynamic address translation protection, page protection, orlow address protection; detecting, based on the attempting access to thestorage location, that the attempted access of the storage locationfailed due to a fault of a storage protection of the at least one typeof storage protection protecting the storage location, wherein thestorage protection is a host level of protection or a guest level ofprotection; determining, based on one or more checks, whether thedetected failure of the attempted access is due to the host level ofprotection or the guest level of protection; and performing an actionbased on the determining, wherein the action performed is one actionbased on the detected failure of the attempted access being due to theguest level of protection and is another action based on the detectedfailure of the attempted access being due to the host level ofprotection.
 2. The computer program product of claim 1, wherein themethod further comprises obtaining information relating to the detectedfailure of the attempted access, the information comprising a hostaddress or a guest address depending on whether the determiningindicates the detected failure of the attempted access is due to thehost level of protection or the guest level of protection, respectively.3. The computer program product of claim 1, wherein the attemptingaccess comprises issuing by a pageable guest of the computingenvironment an instruction that attempts to perform a store into thestorage location.
 4. The computer program product of claim 1, whereinthe method further comprises: performing a guest interrupt, based on thedetermining indicating that the detected failure of the attempted accessis due to the guest level of protection; and performing a hostinterrupt, based on the determining indicating that the detected failureof the attempted access is not due to the guest level of protection. 5.The computer program product of claim 1, wherein the guest level ofprotection is guest dynamic address translation (DAT) protection, andthe host level of protection is host DAT protection.
 6. The computerprogram product of claim 1, wherein the one action comprises indicatingstoring is not permitted by the guest level of protection.
 7. Thecomputer program product of claim 1, wherein the storage location is atleast part of an area of storage that specifies a first frame of hostmain storage, and wherein the other action comprises: based ondetermining that the area of storage is part of a host copy-on-writescheme, assigning a second frame to the area of storage, copyingcontents of the first frame to the second frame, and permitting gueststoring to the area of storage comprising the second frame; and based ondetermining that the area of storage is not part of the hostcopy-on-write scheme, indicating storing is not permitted by the hostlevel of protection.
 8. The computer program product of claim 1, whereinthe failure of the attempted access is detected during execution of apageable guest of the computing environment attempting access to thestorage location.
 9. The computer program product of claim 8, whereinthe pageable guest is of one architecture and is executed by a host ofanother architecture different from the one architecture.
 10. A computersystem for facilitating management of storage of a computingenvironment, the computer system comprising: a memory; and a processorin communication with the memory, wherein the computer system isconfigured to perform a method, said method comprising: attemptingaccess to a storage location, the storage location being protected by atleast one type of storage protection of one or more types of storageprotection, the one or more types of storage protection including atleast one of key controlled protection, access list controlledprotection, dynamic address translation protection, page protection, orlow address protection; detecting, based on the attempting access to thestorage location, that the attempted access of the storage locationfailed due to a fault of a storage protection of the at least one typeof storage protection protecting the storage location, wherein thestorage protection is a host level of protection or a guest level ofprotection; determining, based on one or more checks, whether thedetected failure of the attempted access is due to the host level ofprotection or the guest level of protection; and performing an actionbased on the determining, wherein the action performed is one actionbased on the detected failure of the attempted access being due to theguest level of protection and is another action based on the detectedfailure of the attempted access being due to the host level ofprotection.
 11. The computer system of claim 10, wherein the failure ofthe attempted access is detected during execution of a pageable guest ofthe computing environment attempting access to the storage location. 12.The computer system of claim 10, wherein the method further comprisesobtaining information relating to the detected failure of the attemptedaccess, the information comprising a host address or a guest addressdepending on whether the determining indicates the detected failure ofthe attempted access is due to the host level of protection or the guestlevel of protection, respectively.
 13. The computer system of claim 10,wherein the method further comprises: performing a guest interrupt,based on the determining indicating that the detected failure of theattempted access is due to the guest level of protection; and performinga host interrupt, based on the determining indicating that the detectedfailure of the attempted access is not due to the guest level ofprotection.
 14. The computer system of claim 10, wherein the one actioncomprises indicating storing is not permitted by the guest level ofprotection.
 15. The computer system of claim 10, wherein the storagelocation is at least part of an area of storage that specifies a firstframe of host main storage, and wherein the other action comprises:based on determining that the area of storage is part of a hostcopy-on-write scheme, assigning a second frame to the area of storage,copying contents of the first frame to the second frame, and permittingguest storing to the area of storage comprising the second frame; andbased on determining that the area of storage is not part of the hostcopy-on-write scheme, indicating storing is not permitted by the hostlevel of protection.
 16. A computer-implemented method of facilitatingmanagement of storage of a computing environment, thecomputer-implemented method comprising: attempting, by a processor,access to a storage location, the storage location being protected by atleast one type of storage protection of one or more types of storageprotection, the one or more types of storage protection including atleast one of key controlled protection, access list controlledprotection, dynamic address translation protection, page protection orlow address protection; detecting, based on the attempting access to thestorage location, that the attempted access of the storage locationfailed due to a fault of a storage protection of the at least one typeof storage protection protecting the storage location, wherein thestorage protection is a host level of protection or a guest level ofprotection; determining, based on one or more checks, whether thedetected failure of the attempted access is due to the host level ofprotection or the guest level of protection; and performing an actionbased on the determining, wherein the action performed is one actionbased on the detected failure of the attempted access being due to theguest level of protection and is another action based on the detectedfailure of the attempted access being due to the host level ofprotection.
 17. The computer-implemented method of claim 16, furthercomprising obtaining information relating to the detected failure of theattempted access, the information comprising a host address or a guestaddress depending on whether the determining indicates the detectedfailure of the attempted access is due to the host level of protectionor the guest level of protection, respectively.
 18. Thecomputer-implemented method of claim 16, further comprising: performinga guest interrupt, based on the determining indicating that the detectedfailure of the attempted access is due to the guest level of protection;and performing a host interrupt, based on the determining indicatingthat the detected failure of the attempted access is not due to theguest level of protection.
 19. The computer-implemented method of claim16, wherein the one action comprises indicating storing is not permittedby the guest level of protection.
 20. The computer-implemented method ofclaim 16, wherein the storage location is at least part of an area ofstorage that specifies a first frame of host main storage, and whereinthe other action comprises: based on determining that the area ofstorage is part of a host copy-on-write scheme, assigning a second frameto the area of storage, copying contents of the first frame to thesecond frame, and permitting guest storing to the area of storagecomprising the second frame; and based on determining that the area ofstorage is not part of the host copy-on-write scheme, indicating storingis not permitted by the host level of protection.